# Documentation for the Idris Language¶

Note

The documentation for Idris has been published under the Creative
Commons CC0 License. As such to the extent possible under law, *The
Idris Community* has waived all copyright and related or neighboring
rights to Documentation for Idris.

More information concerning the CC0 can be found online at: http://creativecommons.org/publicdomain/zero/1.0/

## The Idris Tutorial¶

This is the Idris Tutorial. It provides a brief introduction to programming in the Idris Language. It covers the core language features, and assumes some familiarity with an existing functional programming language such as Haskell or OCaml.

Note

The documentation for Idris has been published under the Creative
Commons CC0 License. As such to the extent possible under law, *The
Idris Community* has waived all copyright and related or neighboring
rights to Documentation for Idris.

More information concerning the CC0 can be found online at: http://creativecommons.org/publicdomain/zero/1.0/

### Introduction¶

In conventional programming languages, there is a clear distinction
between *types* and *values*. For example, in Haskell, the following are types, representing
integers, characters, lists of characters, and lists of any value
respectively:

`Int`

,`Char`

,`[Char]`

,`[a]`

Correspondingly, the following values are examples of inhabitants of those types:

`42`

,`’a’`

,`"Hello world!"`

,`[2,3,4,5,6]`

In a language with *dependent types*, however, the distinction is less
clear. Dependent types allow types to “depend” on values — in other
words, types are a *first class* language construct and can be
manipulated like any other value. The standard example is the type of
lists of a given length [1], `Vect n a`

, where `a`

is the element
type and `n`

is the length of the list and can be an arbitrary term.

When types can contain values, and where those values describe
properties, for example the length of a list, the type of a function
can begin to describe its own properties. Take for example the
concatenation of two lists. This operation has the property that the
resulting list’s length is the sum of the lengths of the two input
lists. We can therefore give the following type to the `app`

function, which concatenates vectors:

```
app : Vect n a -> Vect m a -> Vect (n + m) a
```

This tutorial introduces Idris, a general purpose functional
programming language with dependent types. The goal of the Idris
project is to build a dependently typed language suitable for
verifiable general purpose programming. To this end, Idris is a compiled
language which aims to generate efficient executable code. It also has
a lightweight foreign function interface which allows easy interaction
with external `C`

libraries.

#### Intended Audience¶

This tutorial is intended as a brief introduction to the language, and is aimed at readers already familiar with a functional language such as Haskell or OCaml. In particular, a certain amount of familiarity with Haskell syntax is assumed, although most concepts will at least be explained briefly. The reader is also assumed to have some interest in using dependent types for writing and verifying systems software.

For a more in-depth introduction to Idris, which proceeds at a much slower pace, covering interactive program development, with many more examples, see Type-Driven Development with Idris by Edwin Brady, available from Manning.

#### Example Code¶

This tutorial includes some example code, which has been tested with
against Idris. These files are available with the Idris distribution,
so that you can try them out easily. They can be found under
`samples`

. It is, however, strongly recommended that you type
them in yourself, rather than simply loading and reading them.

[1] | Typically, and perhaps confusingly, referred to in the dependently typed programming literature as “vectors” |

### Getting Started¶

#### Prerequisites¶

Before installing Idris, you will need to make sure you have all of the necessary libraries and tools. You will need:

- A fairly recent version of GHC. The earliest version we currently test with is 7.10.3.
- The
*GNU Multiple Precision Arithmetic Library*(GMP) is available from MacPorts/Homebrew and all major Linux distributions.

#### Downloading and Installing¶

The easiest way to install Idris, if you have all of the prerequisites, is to type:

```
cabal update; cabal install idris
```

This will install the latest version released on Hackage, along with any dependencies. If, however, you would like the most up to date development version you can find it, as well as build instructions, on GitHub at: https://github.com/idris-lang/Idris-dev.

If you haven’t previously installed anything using Cabal, then Idris
may not be on your path. Should the Idris executable not be found
please ensure that you have added `~/.cabal/bin`

to your `$PATH`

environment variable. Mac OS X users may find they need to add
`~/Library/Haskell/bin`

instead, and Windows users will typically
find that Cabal installs programs in `%HOME%\AppData\Roaming\cabal\bin`

.

To check that installation has succeeded, and to write your first
Idris program, create a file called `hello.idr`

containing the
following text:

```
module Main
main : IO ()
main = putStrLn "Hello world"
```

If you are familiar with Haskell, it should be fairly clear what the
program is doing and how it works, but if not, we will explain the
details later. You can compile the program to an executable by
entering `idris hello.idr -o hello`

at the shell prompt. This will
create an executable called `hello`

, which you can run:

```
$ idris hello.idr -o hello
$ ./hello
Hello world
```

Please note that the dollar sign `$`

indicates the shell prompt!
Some useful options to the Idris command are:

`-o prog`

to compile to an executable called`prog`

.`--check`

type check the file and its dependencies without starting the interactive environment.`--package pkg`

add package as dependency, e.g.`--package contrib`

to make use of the contrib package.`--help`

display usage summary and command line options.

#### The Interactive Environment¶

Entering `idris`

at the shell prompt starts up the interactive
environment. You should see something like the following:

```
$ idris
____ __ _
/ _/___/ /____(_)____
/ // __ / ___/ / ___/ Version 1.3.0
_/ // /_/ / / / (__ ) http://www.idris-lang.org/
/___/\__,_/_/ /_/____/ Type :? for help
Idris>
```

This gives a `ghci`

style interface which allows evaluation of, as
well as type checking of, expressions; theorem proving, compilation;
editing; and various other operations. The command `:?`

gives a list
of supported commands. Below, we see an example run in
which `hello.idr`

is loaded, the type of `main`

is checked and
then the program is compiled to the executable `hello`

. Type
checking a file, if successful, creates a bytecode version of the file
(in this case `hello.ibc`

) to speed up loading in future. The
bytecode is regenerated if the source file changes.

```
$ idris hello.idr
____ __ _
/ _/___/ /____(_)____
/ // __ / ___/ / ___/ Version 1.3.0
_/ // /_/ / / / (__ ) http://www.idris-lang.org/
/___/\__,_/_/ /_/____/ Type :? for help
Type checking ./hello.idr
*hello> :t main
Main.main : IO ()
*hello> :c hello
*hello> :q
Bye bye
$ ./hello
Hello world
```

### Types and Functions¶

#### Primitive Types¶

Idris defines several primitive types: `Int`

, `Integer`

and
`Double`

for numeric operations, `Char`

and `String`

for text
manipulation, and `Ptr`

which represents foreign pointers. There are
also several data types declared in the library, including `Bool`

,
with values `True`

and `False`

. We can declare some constants with
these types. Enter the following into a file `Prims.idr`

and load it
into the Idris interactive environment by typing `idris Prims.idr`

:

```
module Prims
x : Int
x = 42
foo : String
foo = "Sausage machine"
bar : Char
bar = 'Z'
quux : Bool
quux = False
```

An Idris file consists of an optional module declaration (here
`module Prims`

) followed by an optional list of imports and a
collection of declarations and definitions. In this example no imports
have been specified. However Idris programs can consist of several
modules and the definitions in each module each have their own
namespace. This is discussed further in Section
Modules and Namespaces. When writing Idris programs both the order in which
definitions are given and indentation are significant. Functions and
data types must be defined before use, incidentally each definition must
have a type declaration, for example see `x : Int`

, ```
foo :
String
```

, from the above listing. New declarations must begin at the
same level of indentation as the preceding declaration.
Alternatively, a semicolon `;`

can be used to terminate declarations.

A library module `prelude`

is automatically imported by every
Idris program, including facilities for IO, arithmetic, data
structures and various common functions. The prelude defines several
arithmetic and comparison operators, which we can use at the prompt.
Evaluating things at the prompt gives an answer, and the type of the
answer. For example:

```
*prims> 6*6+6
42 : Integer
*prims> x == 6*6+6
True : Bool
```

All of the usual arithmetic and comparison operators are defined for
the primitive types. They are overloaded using interfaces, as we
will discuss in Section Interfaces and can be extended to
work on user defined types. Boolean expressions can be tested with the
`if...then...else`

construct, for example:

```
*prims> if x == 6 * 6 + 6 then "The answer!" else "Not the answer"
"The answer!" : String
```

#### Data Types¶

Data types are declared in a similar way and with similar syntax to Haskell. Natural numbers and lists, for example, can be declared as follows:

```
data Nat = Z | S Nat -- Natural numbers
-- (zero and successor)
data List a = Nil | (::) a (List a) -- Polymorphic lists
```

The above declarations are taken from the standard library. Unary
natural numbers can be either zero (`Z`

), or the successor of
another natural number (`S k`

). Lists can either be empty (`Nil`

)
or a value added to the front of another list (`x :: xs`

). In the
declaration for `List`

, we used an infix operator `::`

. New
operators such as this can be added using a fixity declaration, as
follows:

```
infixr 10 ::
```

Functions, data constructors and type constructors may all be given
infix operators as names. They may be used in prefix form if enclosed
in brackets, e.g. `(::)`

. Infix operators can use any of the
symbols:

```
:+-*\/=.?|&><!@$%^~#
```

Some operators built from these symbols can’t be user defined. These are
`:`

, `=>`

, `->`

, `<-`

, `=`

, `?=`

, `|`

, `**`

,
`==>`

, `\`

, `%`

, `~`

, `?`

, and `!`

.

#### Functions¶

Functions are implemented by pattern matching, again using a similar
syntax to Haskell. The main difference is that Idris requires type
declarations for all functions, using a single colon `:`

(rather
than Haskell’s double colon `::`

). Some natural number arithmetic
functions can be defined as follows, again taken from the standard
library:

```
-- Unary addition
plus : Nat -> Nat -> Nat
plus Z y = y
plus (S k) y = S (plus k y)
-- Unary multiplication
mult : Nat -> Nat -> Nat
mult Z y = Z
mult (S k) y = plus y (mult k y)
```

The standard arithmetic operators `+`

and `*`

are also overloaded
for use by `Nat`

, and are implemented using the above functions.
Unlike Haskell, there is no restriction on whether types and function
names must begin with a capital letter or not. Function names
(`plus`

and `mult`

above), data constructors (`Z`

, `S`

,
`Nil`

and `::`

) and type constructors (`Nat`

and `List`

) are
all part of the same namespace. By convention, however,
data types and constructor names typically begin with a capital letter.
We can test these functions at the Idris prompt:

```
Idris> plus (S (S Z)) (S (S Z))
4 : Nat
Idris> mult (S (S (S Z))) (plus (S (S Z)) (S (S Z)))
12 : Nat
```

Note

When displaying an element of `Nat`

such as `(S (S (S (S Z))))`

,
Idris displays it as `4`

.
The result of `plus (S (S Z)) (S (S Z))`

is actually `(S (S (S (S Z))))`

which is the natural number `4`

.
This can be checked at the Idris prompt:

```
Idris> (S (S (S (S Z))))
4 : Nat
```

Like arithmetic operations, integer literals are also overloaded using interfaces, meaning that we can also test the functions as follows:

```
Idris> plus 2 2
4 : Nat
Idris> mult 3 (plus 2 2)
12 : Nat
```

You may wonder, by the way, why we have unary natural numbers when our
computers have perfectly good integer arithmetic built in. The reason
is primarily that unary numbers have a very convenient structure which
is easy to reason about, and easy to relate to other data structures
as we will see later. Nevertheless, we do not want this convenience to
be at the expense of efficiency. Fortunately, Idris knows about
the relationship between `Nat`

(and similarly structured types) and
numbers. This means it can optimise the representation, and functions
such as `plus`

and `mult`

.

`where`

clauses¶

Functions can also be defined *locally* using `where`

clauses. For
example, to define a function which reverses a list, we can use an
auxiliary function which accumulates the new, reversed list, and which
does not need to be visible globally:

```
reverse : List a -> List a
reverse xs = revAcc [] xs where
revAcc : List a -> List a -> List a
revAcc acc [] = acc
revAcc acc (x :: xs) = revAcc (x :: acc) xs
```

Indentation is significant — functions in the `where`

block must be
indented further than the outer function.

Note

Scope

Any names which are visible in the outer scope are also visible in
the `where`

clause (unless they have been redefined, such as `xs`

here). A name which appears only in the type will be in scope in the
`where`

clause if it is a *parameter* to one of the types, i.e. it
is fixed across the entire structure.

As well as functions, `where`

blocks can include local data
declarations, such as the following where `MyLT`

is not accessible
outside the definition of `foo`

:

```
foo : Int -> Int
foo x = case isLT of
Yes => x*2
No => x*4
where
data MyLT = Yes | No
isLT : MyLT
isLT = if x < 20 then Yes else No
```

In general, functions defined in a `where`

clause need a type
declaration just like any top level function. However, the type
declaration for a function `f`

*can* be omitted if:

`f`

appears in the right hand side of the top level definition- The type of
`f`

can be completely determined from its first application

So, for example, the following definitions are legal:

```
even : Nat -> Bool
even Z = True
even (S k) = odd k where
odd Z = False
odd (S k) = even k
test : List Nat
test = [c (S 1), c Z, d (S Z)]
where c x = 42 + x
d y = c (y + 1 + z y)
where z w = y + w
```

##### Holes¶

Idris programs can contain *holes* which stand for incomplete parts of
programs. For example, we could leave a hole for the greeting in our
“Hello world” program:

```
main : IO ()
main = putStrLn ?greeting
```

The syntax `?greeting`

introduces a hole, which stands for a part of
a program which is not yet written. This is a valid Idris program, and you
can check the type of `greeting`

:

```
*Hello> :t greeting
--------------------------------------
greeting : String
```

Checking the type of a hole also shows the types of any variables in scope.
For example, given an incomplete definition of `even`

:

```
even : Nat -> Bool
even Z = True
even (S k) = ?even_rhs
```

We can check the type of `even_rhs`

and see the expected return type,
and the type of the variable `k`

:

```
*Even> :t even_rhs
k : Nat
--------------------------------------
even_rhs : Bool
```

Holes are useful because they help us write functions *incrementally*.
Rather than writing an entire function in one go, we can leave some parts
unwritten and use Idris to tell us what is necessary to complete the
definition.

#### Dependent Types¶

##### First Class Types¶

In Idris, types are first class, meaning that they can be computed and manipulated (and passed to functions) just like any other language construct. For example, we could write a function which computes a type:

```
isSingleton : Bool -> Type
isSingleton True = Nat
isSingleton False = List Nat
```

This function calculates the appropriate type from a `Bool`

which flags
whether the type should be a singleton or not. We can use this function
to calculate a type anywhere that a type can be used. For example, it
can be used to calculate a return type:

```
mkSingle : (x : Bool) -> isSingleton x
mkSingle True = 0
mkSingle False = []
```

Or it can be used to have varying input types. The following function
calculates either the sum of a list of `Nat`

, or returns the given
`Nat`

, depending on whether the singleton flag is true:

```
sum : (single : Bool) -> isSingleton single -> Nat
sum True x = x
sum False [] = 0
sum False (x :: xs) = x + sum False xs
```

##### Vectors¶

A standard example of a dependent data type is the type of “lists with
length”, conventionally called vectors in the dependent type
literature. They are available as part of the Idris library, by
importing `Data.Vect`

, or we can declare them as follows:

```
data Vect : Nat -> Type -> Type where
Nil : Vect Z a
(::) : a -> Vect k a -> Vect (S k) a
```

Note that we have used the same constructor names as for `List`

.
Ad-hoc name overloading such as this is accepted by Idris,
provided that the names are declared in different namespaces (in
practice, normally in different modules). Ambiguous constructor names
can normally be resolved from context.

This declares a family of types, and so the form of the declaration is
rather different from the simple type declarations above. We
explicitly state the type of the type constructor `Vect`

— it takes
a `Nat`

and a type as an argument, where `Type`

stands for the
type of types. We say that `Vect`

is *indexed* over `Nat`

and
*parameterised* by `Type`

. Each constructor targets a different part
of the family of types. `Nil`

can only be used to construct vectors
with zero length, and `::`

to construct vectors with non-zero
length. In the type of `::`

, we state explicitly that an element of
type `a`

and a tail of type `Vect k a`

(i.e., a vector of length
`k`

) combine to make a vector of length `S k`

.

We can define functions on dependent types such as `Vect`

in the same
way as on simple types such as `List`

and `Nat`

above, by pattern
matching. The type of a function over `Vect`

will describe what
happens to the lengths of the vectors involved. For example, `++`

,
defined as follows, appends two `Vect`

:

```
(++) : Vect n a -> Vect m a -> Vect (n + m) a
(++) Nil ys = ys
(++) (x :: xs) ys = x :: xs ++ ys
```

The type of `(++)`

states that the resulting vector’s length will be
the sum of the input lengths. If we get the definition wrong in such a
way that this does not hold, Idris will not accept the definition.
For example:

```
(++) : Vect n a -> Vect m a -> Vect (n + m) a
(++) Nil ys = ys
(++) (x :: xs) ys = x :: xs ++ xs -- BROKEN
```

When run through the Idris type checker, this results in the following:

```
$ idris VBroken.idr --check
VBroken.idr:9:23-25:
When checking right hand side of Vect.++ with expected type
Vect (S k + m) a
When checking an application of constructor Vect.:::
Type mismatch between
Vect (k + k) a (Type of xs ++ xs)
and
Vect (plus k m) a (Expected type)
Specifically:
Type mismatch between
plus k k
and
plus k m
```

This error message suggests that there is a length mismatch between
two vectors — we needed a vector of length `k + m`

, but provided a
vector of length `k + k`

.

##### The Finite Sets¶

Finite sets, as the name suggests, are sets with a finite number of
elements. They are available as part of the Idris library, by
importing `Data.Fin`

, or can be declared as follows:

```
data Fin : Nat -> Type where
FZ : Fin (S k)
FS : Fin k -> Fin (S k)
```

From the signature, we can see that this is a type constructor that takes a `Nat`

, and produces a type.
So this is not a set in the sense of a collection that is a container of objects,
rather it is the canonical set of unnamed elements, as in “the set of 5 elements,” for example.
Effectively, it is a type that captures integers that fall into the range of zero to `(n - 1)`

where
`n`

is the argument used to instantiate the `Fin`

type.
For example, `Fin 5`

can be thought of as the type of integers between 0 and 4.

Let us look at the constructors in greater detail.

`FZ`

is the zeroth element of a finite set with `S k`

elements;
`FS n`

is the `n+1`

th element of a finite set with `S k`

elements. `Fin`

is indexed by a `Nat`

, which represents the number
of elements in the set. Since we can’t construct an element of an
empty set, neither constructor targets `Fin Z`

.

As mentioned above, a useful application of the `Fin`

family is to
represent bounded natural numbers. Since the first `n`

natural
numbers form a finite set of `n`

elements, we can treat `Fin n`

as
the set of integers greater than or equal to zero and less than `n`

.

For example, the following function which looks up an element in a
`Vect`

, by a bounded index given as a `Fin n`

, is defined in the
prelude:

```
index : Fin n -> Vect n a -> a
index FZ (x :: xs) = x
index (FS k) (x :: xs) = index k xs
```

This function looks up a value at a given location in a vector. The
location is bounded by the length of the vector (`n`

in each case),
so there is no need for a run-time bounds check. The type checker
guarantees that the location is no larger than the length of the
vector, and of course no less than zero.

Note also that there is no case for `Nil`

here. This is because it
is impossible. Since there is no element of `Fin Z`

, and the
location is a `Fin n`

, then `n`

can not be `Z`

. As a result,
attempting to look up an element in an empty vector would give a
compile time type error, since it would force `n`

to be `Z`

.

##### Implicit Arguments¶

Let us take a closer look at the type of `index`

:

```
index : Fin n -> Vect n a -> a
```

It takes two arguments, an element of the finite set of `n`

elements,
and a vector with `n`

elements of type `a`

. But there are also two
names, `n`

and `a`

, which are not declared explicitly. These are
*implicit* arguments to `index`

. We could also write the type of
`index`

as:

```
index : {a:Type} -> {n:Nat} -> Fin n -> Vect n a -> a
```

Implicit arguments, given in braces `{}`

in the type declaration,
are not given in applications of `index`

; their values can be
inferred from the types of the `Fin n`

and `Vect n a`

arguments. Any name beginning with a lower case letter which appears
as a parameter or index in a
type declaration, which is not applied to any arguments, will
*always* be automatically
bound as an implicit argument. Implicit arguments can still be given
explicitly in applications, using `{a=value}`

and `{n=value}`

, for
example:

```
index {a=Int} {n=2} FZ (2 :: 3 :: Nil)
```

In fact, any argument, implicit or explicit, may be given a name. We
could have declared the type of `index`

as:

```
index : (i:Fin n) -> (xs:Vect n a) -> a
```

It is a matter of taste whether you want to do this — sometimes it can help document a function by making the purpose of an argument more clear.

Furthermore, `{}`

can be used to pattern match on the left hand side, i.e.
`{var = pat}`

gets an implicit variable and attempts to pattern match on “pat”;
For example:

```
isEmpty : Vect n a -> Bool
isEmpty {n = Z} _ = True
isEmpty {n = S k} _ = False
```

##### “`using`

” notation¶

Sometimes it is useful to provide types of implicit arguments,
particularly where there is a dependency ordering, or where the
implicit arguments themselves have dependencies. For example, we may
wish to state the types of the implicit arguments in the following
definition, which defines a predicate on vectors (this is also defined
in `Data.Vect`

, under the name `Elem`

):

```
data IsElem : a -> Vect n a -> Type where
Here : {x:a} -> {xs:Vect n a} -> IsElem x (x :: xs)
There : {x,y:a} -> {xs:Vect n a} -> IsElem x xs -> IsElem x (y :: xs)
```

An instance of `IsElem x xs`

states that `x`

is an element of
`xs`

. We can construct such a predicate if the required element is
`Here`

, at the head of the vector, or `There`

, in the tail of the
vector. For example:

```
testVec : Vect 4 Int
testVec = 3 :: 4 :: 5 :: 6 :: Nil
inVect : IsElem 5 Main.testVec
inVect = There (There Here)
```

Important

Implicit Arguments and Scope

Within the type signature the typechecker will treat all variables
that start with an lowercase letter **and** are not applied to
something else as an implicit variable. To get the above code
example to compile you will need to provide a qualified name for
`testVec`

. In the example above, we have assumed that the code
lives within the `Main`

module.

If the same implicit arguments are being used a lot, it can make a
definition difficult to read. To avoid this problem, a `using`

block
gives the types and ordering of any implicit arguments which can
appear within the block:

```
using (x:a, y:a, xs:Vect n a)
data IsElem : a -> Vect n a -> Type where
Here : IsElem x (x :: xs)
There : IsElem x xs -> IsElem x (y :: xs)
```

###### Note: Declaration Order and `mutual`

blocks¶

In general, functions and data types must be defined before use, since
dependent types allow functions to appear as part of types, and type
checking can rely on how particular functions are defined (though this
is only true of total functions; see Section Totality Checking).
However, this restriction can be relaxed by using a `mutual`

block,
which allows data types and functions to be defined simultaneously:

```
mutual
even : Nat -> Bool
even Z = True
even (S k) = odd k
odd : Nat -> Bool
odd Z = False
odd (S k) = even k
```

In a `mutual`

block, first all of the type declarations are added,
then the function bodies. As a result, none of the function types can
depend on the reduction behaviour of any of the functions in the
block.

#### I/O¶

Computer programs are of little use if they do not interact with the
user or the system in some way. The difficulty in a pure language such
as Idris — that is, a language where expressions do not have
side-effects — is that I/O is inherently side-effecting. Therefore in
Idris, such interactions are encapsulated in the type `IO`

:

```
data IO a -- IO operation returning a value of type a
```

We’ll leave the definition of `IO`

abstract, but effectively it
describes what the I/O operations to be executed are, rather than how
to execute them. The resulting operations are executed externally, by
the run-time system. We’ve already seen one IO program:

```
main : IO ()
main = putStrLn "Hello world"
```

The type of `putStrLn`

explains that it takes a string, and returns
an element of the unit type `()`

via an I/O action. There is a
variant `putStr`

which outputs a string without a newline:

```
putStrLn : String -> IO ()
putStr : String -> IO ()
```

We can also read strings from user input:

```
getLine : IO String
```

A number of other I/O operations are defined in the prelude, for example for reading and writing files, including:

```
data File -- abstract
data Mode = Read | Write | ReadWrite
openFile : (f : String) -> (m : Mode) -> IO (Either FileError File)
closeFile : File -> IO ()
fGetLine : (h : File) -> IO (Either FileError String)
fPutStr : (h : File) -> (str : String) -> IO (Either FileError ())
fEOF : File -> IO Bool
```

Note that several of these return `Either`

, since they may fail.

#### “`do`

” notation¶

I/O programs will typically need to sequence actions, feeding the
output of one computation into the input of the next. `IO`

is an
abstract type, however, so we can’t access the result of a computation
directly. Instead, we sequence operations with `do`

notation:

```
greet : IO ()
greet = do putStr "What is your name? "
name <- getLine
putStrLn ("Hello " ++ name)
```

The syntax `x <- iovalue`

executes the I/O operation `iovalue`

, of
type `IO a`

, and puts the result, of type `a`

into the variable
`x`

. In this case, `getLine`

returns an `IO String`

, so `name`

has type `String`

. Indentation is significant — each statement in
the do block must begin in the same column. The `pure`

operation
allows us to inject a value directly into an IO operation:

```
pure : a -> IO a
```

As we will see later, `do`

notation is more general than this, and
can be overloaded.

#### Laziness¶

Normally, arguments to functions are evaluated before the function
itself (that is, Idris uses *eager* evaluation). However, this is
not always the best approach. Consider the following function:

```
ifThenElse : Bool -> a -> a -> a
ifThenElse True t e = t
ifThenElse False t e = e
```

This function uses one of the `t`

or `e`

arguments, but not both
(in fact, this is used to implement the `if...then...else`

construct
as we will see later). We would prefer if *only* the argument which was
used was evaluated. To achieve this, Idris provides a `Lazy`

data type, which allows evaluation to be suspended:

```
data Lazy : Type -> Type where
Delay : (val : a) -> Lazy a
Force : Lazy a -> a
```

A value of type `Lazy a`

is unevaluated until it is forced by
`Force`

. The Idris type checker knows about the `Lazy`

type,
and inserts conversions where necessary between `Lazy a`

and `a`

,
and vice versa. We can therefore write `ifThenElse`

as follows,
without any explicit use of `Force`

or `Delay`

:

```
ifThenElse : Bool -> Lazy a -> Lazy a -> a
ifThenElse True t e = t
ifThenElse False t e = e
```

#### Codata Types¶

Codata types allow us to define infinite data structures by marking recursive
arguments as potentially infinite. For
a codata type `T`

, each of its constructor arguments of type `T`

are transformed
into an argument of type `Inf T`

. This makes each of the `T`

arguments
lazy, and allows infinite data structures of type `T`

to be built. One
example of a codata type is Stream, which is defined as follows.

```
codata Stream : Type -> Type where
(::) : (e : a) -> Stream a -> Stream a
```

This gets translated into the following by the compiler.

```
data Stream : Type -> Type where
(::) : (e : a) -> Inf (Stream a) -> Stream a
```

The following is an example of how the codata type `Stream`

can be used to
form an infinite data structure. In this case we are creating an infinite stream
of ones.

```
ones : Stream Nat
ones = 1 :: ones
```

It is important to note that codata does not allow the creation of infinite mutually recursive data structures. For example the following will create an infinite loop and cause a stack overflow.

```
mutual
codata Blue a = B a (Red a)
codata Red a = R a (Blue a)
mutual
blue : Blue Nat
blue = B 1 red
red : Red Nat
red = R 1 blue
mutual
findB : (a -> Bool) -> Blue a -> a
findB f (B x r) = if f x then x else findR f r
findR : (a -> Bool) -> Red a -> a
findR f (R x b) = if f x then x else findB f b
main : IO ()
main = do printLn $ findB (== 1) blue
```

To fix this we must add explicit `Inf`

declarations to the constructor
parameter types, since codata will not add it to constructor parameters of a
**different** type from the one being defined. For example, the following
outputs `1`

.

```
mutual
data Blue : Type -> Type where
B : a -> Inf (Red a) -> Blue a
data Red : Type -> Type where
R : a -> Inf (Blue a) -> Red a
mutual
blue : Blue Nat
blue = B 1 red
red : Red Nat
red = R 1 blue
mutual
findB : (a -> Bool) -> Blue a -> a
findB f (B x r) = if f x then x else findR f r
findR : (a -> Bool) -> Red a -> a
findR f (R x b) = if f x then x else findB f b
main : IO ()
main = do printLn $ findB (== 1) blue
```

#### Useful Data Types¶

Idris includes a number of useful data types and library functions
(see the `libs/`

directory in the distribution, and the
documentation). This section
describes a few of these. The functions described here are imported
automatically by every Idris program, as part of `Prelude.idr`

.

`List`

and `Vect`

¶

We have already seen the `List`

and `Vect`

data types:

```
data List a = Nil | (::) a (List a)
data Vect : Nat -> Type -> Type where
Nil : Vect Z a
(::) : a -> Vect k a -> Vect (S k) a
```

Note that the constructor names are the same for each — constructor
names (in fact, names in general) can be overloaded, provided that
they are declared in different namespaces (see Section
Modules and Namespaces), and will typically be resolved according to
their type. As syntactic sugar, any type with the constructor names
`Nil`

and `::`

can be written in list form. For example:

`[]`

means`Nil`

`[1,2,3]`

means`1 :: 2 :: 3 :: Nil`

The library also defines a number of functions for manipulating these
types. `map`

is overloaded both for `List`

and `Vect`

and
applies a function to every element of the list or vector.

```
map : (a -> b) -> List a -> List b
map f [] = []
map f (x :: xs) = f x :: map f xs
map : (a -> b) -> Vect n a -> Vect n b
map f [] = []
map f (x :: xs) = f x :: map f xs
```

For example, given the following vector of integers, and a function to double an integer:

```
intVec : Vect 5 Int
intVec = [1, 2, 3, 4, 5]
double : Int -> Int
double x = x * 2
```

the function `map`

can be used as follows to double every element in
the vector:

```
*UsefulTypes> show (map double intVec)
"[2, 4, 6, 8, 10]" : String
```

For more details of the functions available on `List`

and
`Vect`

, look in the library files:

`libs/prelude/Prelude/List.idr`

`libs/base/Data/List.idr`

`libs/base/Data/Vect.idr`

`libs/base/Data/VectType.idr`

Functions include filtering, appending, reversing, and so on.

###### Aside: Anonymous functions and operator sections¶

There are actually neater ways to write the above expression. One way would be to use an anonymous function:

```
*UsefulTypes> show (map (\x => x * 2) intVec)
"[2, 4, 6, 8, 10]" : String
```

The notation `\x => val`

constructs an anonymous function which takes
one argument, `x`

and returns the expression `val`

. Anonymous
functions may take several arguments, separated by commas,
e.g. `\x, y, z => val`

. Arguments may also be given explicit types,
e.g. `\x : Int => x * 2`

, and can pattern match,
e.g. `\(x, y) => x + y`

. We could also use an operator section:

```
*UsefulTypes> show (map (* 2) intVec)
"[2, 4, 6, 8, 10]" : String
```

`(*2)`

is shorthand for a function which multiplies a number
by 2. It expands to `\x => x * 2`

. Similarly, `(2*)`

would expand
to `\x => 2 * x`

.

##### Maybe¶

`Maybe`

describes an optional value. Either there is a value of the
given type, or there isn’t:

```
data Maybe a = Just a | Nothing
```

`Maybe`

is one way of giving a type to an operation that may
fail. For example, looking something up in a `List`

(rather than a
vector) may result in an out of bounds error:

```
list_lookup : Nat -> List a -> Maybe a
list_lookup _ Nil = Nothing
list_lookup Z (x :: xs) = Just x
list_lookup (S k) (x :: xs) = list_lookup k xs
```

The `maybe`

function is used to process values of type `Maybe`

,
either by applying a function to the value, if there is one, or by
providing a default value:

```
maybe : Lazy b -> Lazy (a -> b) -> Maybe a -> b
```

Note that the types of the first two arguments are wrapped in
`Lazy`

. Since only one of the two arguments will actually be used,
we mark them as `Lazy`

in case they are large expressions where it
would be wasteful to compute and then discard them.

##### Tuples¶

Values can be paired with the following built-in data type:

```
data Pair a b = MkPair a b
```

As syntactic sugar, we can write `(a, b)`

which, according to
context, means either `Pair a b`

or `MkPair a b`

. Tuples can
contain an arbitrary number of values, represented as nested pairs:

```
fred : (String, Int)
fred = ("Fred", 42)
jim : (String, Int, String)
jim = ("Jim", 25, "Cambridge")
```

```
*UsefulTypes> fst jim
"Jim" : String
*UsefulTypes> snd jim
(25, "Cambridge") : (Int, String)
*UsefulTypes> jim == ("Jim", (25, "Cambridge"))
True : Bool
```

##### Dependent Pairs¶

Dependent pairs allow the type of the second element of a pair to depend on the value of the first element:

```
data DPair : (a : Type) -> (P : a -> Type) -> Type where
MkDPair : {P : a -> Type} -> (x : a) -> P x -> DPair a P
```

Again, there is syntactic sugar for this. `(a : A ** P)`

is the type
of a pair of A and P, where the name `a`

can occur inside `P`

.
`( a ** p )`

constructs a value of this type. For example, we can
pair a number with a `Vect`

of a particular length:

```
vec : (n : Nat ** Vect n Int)
vec = (2 ** [3, 4])
```

If you like, you can write it out the long way, the two are precisely equivalent:

```
vec : DPair Nat (\n => Vect n Int)
vec = MkDPair 2 [3, 4]
```

The type checker could of course infer the value of the first element
from the length of the vector. We can write an underscore `_`

in
place of values which we expect the type checker to fill in, so the
above definition could also be written as:

```
vec : (n : Nat ** Vect n Int)
vec = (_ ** [3, 4])
```

We might also prefer to omit the type of the first element of the pair, since, again, it can be inferred:

```
vec : (n ** Vect n Int)
vec = (_ ** [3, 4])
```

One use for dependent pairs is to return values of dependent types
where the index is not necessarily known in advance. For example, if
we filter elements out of a `Vect`

according to some predicate, we
will not know in advance what the length of the resulting vector will
be:

```
filter : (a -> Bool) -> Vect n a -> (p ** Vect p a)
```

If the `Vect`

is empty, the result is easy:

```
filter p Nil = (_ ** [])
```

In the `::`

case, we need to inspect the result of a recursive call
to `filter`

to extract the length and the vector from the result. To
do this, we use `with`

notation, which allows pattern matching on
intermediate values:

```
filter p (x :: xs) with (filter p xs)
| ( _ ** xs' ) = if (p x) then ( _ ** x :: xs' ) else ( _ ** xs' )
```

We will see more on `with`

notation later.

Dependent pairs are sometimes referred to as “Sigma types”.

##### Records¶

*Records* are data types which collect several values (the record’s
*fields*) together. Idris provides syntax for defining records and
automatically generating field access and update functions. Unlike
the syntax used for data structures, records in Idris follow a
different syntax to that seen with Haskell. For example, we can
represent a person’s name and age in a record:

```
record Person where
constructor MkPerson
firstName, middleName, lastName : String
age : Int
fred : Person
fred = MkPerson "Fred" "Joe" "Bloggs" 30
```

The constructor name is provided using the `constructor`

keyword,
and the *fields* are then given which are in an indented block
following the where keyword (here, `firstName`

, `middleName`

,
`lastName`

, and `age`

). You can declare multiple fields on a
single line, provided that they have the same type. The field names
can be used to access the field values:

```
*Record> firstName fred
"Fred" : String
*Record> age fred
30 : Int
*Record> :t firstName
firstName : Person -> String
```

We can also use the field names to update a record (or, more precisely, produce a copy of the record with the given fields updated):

```
*Record> record { firstName = "Jim" } fred
MkPerson "Jim" "Joe" "Bloggs" 30 : Person
*Record> record { firstName = "Jim", age $= (+ 1) } fred
MkPerson "Jim" "Joe" "Bloggs" 31 : Person
```

The syntax `record { field = val, ... }`

generates a function which
updates the given fields in a record. `=`

assigns a new value to a field,
and `$=`

applies a function to update its value.

Each record is defined in its own namespace, which means that field names can be reused in multiple records.

Records, and fields within records, can have dependent types. Updates are allowed to change the type of a field, provided that the result is well-typed.

```
record Class where
constructor ClassInfo
students : Vect n Person
className : String
```

It is safe to update the `students`

field to a vector of a different
length because it will not affect the type of the record:

```
addStudent : Person -> Class -> Class
addStudent p c = record { students = p :: students c } c
```

```
*Record> addStudent fred (ClassInfo [] "CS")
ClassInfo [MkPerson "Fred" "Joe" "Bloggs" 30] "CS" : Class
```

We could also use `$=`

to define `addStudent`

more concisely:

```
addStudent' : Person -> Class -> Class
addStudent' p c = record { students $= (p ::) } c
```

###### Nested record update¶

Idris also provides a convenient syntax for accessing and updating
nested records. For example, if a field is accessible with the
expression `c (b (a x))`

, it can be updated using the following
syntax:

```
record { a->b->c = val } x
```

This returns a new record, with the field accessed by the path
`a->b->c`

set to `val`

. The syntax is first class, i.e. ```
record {
a->b->c = val }
```

itself has a function type. Symmetrically, the field
can also be accessed with the following syntax:

```
record { a->b->c } x
```

The `$=`

notation is also valid for nested record updates.

##### Dependent Records¶

Records can also be dependent on values. Records have *parameters*, which
cannot be updated like the other fields. The parameters appear as arguments
to the resulting type, and are written following the record type
name. For example, a pair type could be defined as follows:

```
record Prod a b where
constructor Times
fst : a
snd : b
```

Using the `class`

record from earlier, the size of the class can be
restricted using a `Vect`

and the size included in the type by parameterising
the record with the size. For example:

```
record SizedClass (size : Nat) where
constructor SizedClassInfo
students : Vect size Person
className : String
```

**Note** that it is no longer possible to use the `addStudent`

function from earlier, since that would change the size of the class. A
function to add a student must now specify in the type that the
size of the class has been increased by one. As the size is specified
using natural numbers, the new value can be incremented using the
`S`

constructor:

```
addStudent : Person -> SizedClass n -> SizedClass (S n)
addStudent p c = SizedClassInfo (p :: students c) (className c)
```

#### More Expressions¶

`let`

bindings¶

Intermediate values can be calculated using `let`

bindings:

```
mirror : List a -> List a
mirror xs = let xs' = reverse xs in
xs ++ xs'
```

We can do simple pattern matching in `let`

bindings too. For
example, we can extract fields from a record as follows, as well as by
pattern matching at the top level:

```
data Person = MkPerson String Int
showPerson : Person -> String
showPerson p = let MkPerson name age = p in
name ++ " is " ++ show age ++ " years old"
```

##### List comprehensions¶

Idris provides *comprehension* notation as a convenient shorthand
for building lists. The general form is:

```
[ expression | qualifiers ]
```

This generates the list of values produced by evaluating the
`expression`

, according to the conditions given by the comma
separated `qualifiers`

. For example, we can build a list of
Pythagorean triples as follows:

```
pythag : Int -> List (Int, Int, Int)
pythag n = [ (x, y, z) | z <- [1..n], y <- [1..z], x <- [1..y],
x*x + y*y == z*z ]
```

The `[a..b]`

notation is another shorthand which builds a list of
numbers between `a`

and `b`

. Alternatively `[a,b..c]`

builds a
list of numbers between `a`

and `c`

with the increment specified
by the difference between `a`

and `b`

. This works for type `Nat`

,
`Int`

and `Integer`

, using the `enumFromTo`

and `enumFromThenTo`

function from the prelude.

`case`

expressions¶

Another way of inspecting intermediate values of *simple* types is to
use a `case`

expression. The following function, for example, splits
a string into two at a given character:

```
splitAt : Char -> String -> (String, String)
splitAt c x = case break (== c) x of
(x, y) => (x, strTail y)
```

`break`

is a library function which breaks a string into a pair of
strings at the point where the given function returns true. We then
deconstruct the pair it returns, and remove the first character of the
second string.

A `case`

expression can match several cases, for example, to inspect
an intermediate value of type `Maybe a`

. Recall `list_lookup`

which looks up an index in a list, returning `Nothing`

if the index
is out of bounds. We can use this to write `lookup_default`

, which
looks up an index and returns a default value if the index is out of
bounds:

```
lookup_default : Nat -> List a -> a -> a
lookup_default i xs def = case list_lookup i xs of
Nothing => def
Just x => x
```

If the index is in bounds, we get the value at that index, otherwise we get a default value:

```
*UsefulTypes> lookup_default 2 [3,4,5,6] (-1)
5 : Integer
*UsefulTypes> lookup_default 4 [3,4,5,6] (-1)
-1 : Integer
```

**Restrictions:** The `case`

construct is intended for simple
analysis of intermediate expressions to avoid the need to write
auxiliary functions, and is also used internally to implement pattern
matching `let`

and lambda bindings. It will *only* work if:

- Each branch
*matches*a value of the same type, and*returns*a value of the same type. - The type of the result is “known”. i.e. the type of the expression
can be determined
*without*type checking the`case`

-expression itself.

#### Totality¶

Idris distinguishes between *total* and *partial* functions.
A total function is a function that either:

- Terminates for all possible inputs, or
- Produces a non-empty, finite, prefix of a possibly infinite result

If a function is total, we can consider its type a precise description of what
that function will do. For example, if we have a function with a return
type of `String`

we know something different, depending on whether or not
it’s total:

- If it’s total, it will return a value of type
`String`

in finite time; - If it’s partial, then as long as it doesn’t crash or enter an infinite loop,
it will return a
`String`

.

Idris makes this distinction so that it knows which functions are safe to evaluate while type checking (as we’ve seen with First Class Types). After all, if it tries to evaluate a function during type checking which doesn’t terminate, then type checking won’t terminate! Therefore, only total functions will be evaluated during type checking. Partial functions can still be used in types, but will not be evaluated further.

### Interfaces¶

We often want to define functions which work across several different
data types. For example, we would like arithmetic operators to work on
`Int`

, `Integer`

and `Double`

at the very least. We would like
`==`

to work on the majority of data types. We would like to be able
to display different types in a uniform way.

To achieve this, we use *interfaces*, which are similar to type classes in
Haskell or traits in Rust. To define an interface, we provide a collection of
overloadable functions. A simple example is the `Show`

interface, which is defined in the prelude and provides an interface for
converting values to `String`

:

```
interface Show a where
show : a -> String
```

This generates a function of the following type (which we call a
*method* of the `Show`

interface):

```
show : Show a => a -> String
```

We can read this as: “under the constraint that `a`

has an implementation
of `Show`

, take an input `a`

and return a `String`

.” An implementation
of an interface is defined by giving definitions of the methods of the interface.
For example, the `Show`

implementation for `Nat`

could be defined as:

```
Show Nat where
show Z = "Z"
show (S k) = "s" ++ show k
```

```
Idris> show (S (S (S Z)))
"sssZ" : String
```

Only one implementation of an interface can be given for a type — implementations may
not overlap. Implementation declarations can themselves have constraints.
To help with resolution, the arguments of an implementation must be
constructors (either data or type constructors) or variables
(i.e. you cannot give an implementation for a function). For
example, to define a `Show`

implementation for vectors, we need to know
that there is a `Show`

implementation for the element type, because we are
going to use it to convert each element to a `String`

:

```
Show a => Show (Vect n a) where
show xs = "[" ++ show' xs ++ "]" where
show' : Vect n a -> String
show' Nil = ""
show' (x :: Nil) = show x
show' (x :: xs) = show x ++ ", " ++ show' xs
```

#### Default Definitions¶

The library defines an `Eq`

interface which provides methods for
comparing values for equality or inequality, with implementations for all of
the built-in types:

```
interface Eq a where
(==) : a -> a -> Bool
(/=) : a -> a -> Bool
```

To declare an implementation for a type, we have to give definitions of all
of the methods. For example, for an implementation of `Eq`

for `Nat`

:

```
Eq Nat where
Z == Z = True
(S x) == (S y) = x == y
Z == (S y) = False
(S x) == Z = False
x /= y = not (x == y)
```

It is hard to imagine many cases where the `/=`

method will be
anything other than the negation of the result of applying the `==`

method. It is therefore convenient to give a default definition for
each method in the interface declaration, in terms of the other method:

```
interface Eq a where
(==) : a -> a -> Bool
(/=) : a -> a -> Bool
x /= y = not (x == y)
x == y = not (x /= y)
```

A minimal complete implementation of `Eq`

requires either
`==`

or `/=`

to be defined, but does not require both. If a method
definition is missing, and there is a default definition for it, then
the default is used instead.

#### Extending Interfaces¶

Interfaces can also be extended. A logical next step from an equality
relation `Eq`

is to define an ordering relation `Ord`

. We can
define an `Ord`

interface which inherits methods from `Eq`

as well as
defining some of its own:

```
data Ordering = LT | EQ | GT
```

```
interface Eq a => Ord a where
compare : a -> a -> Ordering
(<) : a -> a -> Bool
(>) : a -> a -> Bool
(<=) : a -> a -> Bool
(>=) : a -> a -> Bool
max : a -> a -> a
min : a -> a -> a
```

The `Ord`

interface allows us to compare two values and determine their
ordering. Only the `compare`

method is required; every other method
has a default definition. Using this we can write functions such as
`sort`

, a function which sorts a list into increasing order,
provided that the element type of the list is in the `Ord`

interface. We
give the constraints on the type variables left of the fat arrow
`=>`

, and the function type to the right of the fat arrow:

```
sort : Ord a => List a -> List a
```

Functions, interfaces and implementations can have multiple constraints. Multiple constraints are written in brackets in a comma separated list, for example:

```
sortAndShow : (Ord a, Show a) => List a -> String
sortAndShow xs = show (sort xs)
```

##### Note: Interfaces and `mutual`

blocks¶

Idris is strictly “define before use”, except in `mutual`

blocks.
In a `mutual`

block, Idris elaborates in two passes: types on the first
pass and definitions on the second. When the mutual block contains an
interface declaration, it elaborates the interface header but none of the
method types on the first pass, and elaborates the method types and any
default definitions on the second pass.

#### Functors and Applicatives¶

So far, we have seen single parameter interfaces, where the parameter
is of type `Type`

. In general, there can be any number of parameters
(even zero), and the parameters can have *any* type. If the type
of the parameter is not `Type`

, we need to give an explicit type
declaration. For example, the `Functor`

interface is defined in the
prelude:

```
interface Functor (f : Type -> Type) where
map : (m : a -> b) -> f a -> f b
```

A functor allows a function to be applied across a structure, for
example to apply a function to every element in a `List`

:

```
Functor List where
map f [] = []
map f (x::xs) = f x :: map f xs
```

```
Idris> map (*2) [1..10]
[2, 4, 6, 8, 10, 12, 14, 16, 18, 20] : List Integer
```

Having defined `Functor`

, we can define `Applicative`

which
abstracts the notion of function application:

```
infixl 2 <*>
interface Functor f => Applicative (f : Type -> Type) where
pure : a -> f a
(<*>) : f (a -> b) -> f a -> f b
```

#### Monads and `do`

-notation¶

The `Monad`

interface allows us to encapsulate binding and computation,
and is the basis of `do`

-notation introduced in Section
“do” notation. It extends `Applicative`

as defined above, and is
defined as follows:

```
interface Applicative m => Monad (m : Type -> Type) where
(>>=) : m a -> (a -> m b) -> m b
```

Inside a `do`

block, the following syntactic transformations are
applied:

`x <- v; e`

becomes`v >>= (\x => e)`

`v; e`

becomes`v >>= (\_ => e)`

`let x = v; e`

becomes`let x = v in e`

`IO`

has an implementation of `Monad`

, defined using primitive functions.
We can also define an implementation for `Maybe`

, as follows:

```
Monad Maybe where
Nothing >>= k = Nothing
(Just x) >>= k = k x
```

Using this we can, for example, define a function which adds two
`Maybe Int`

, using the monad to encapsulate the error handling:

```
m_add : Maybe Int -> Maybe Int -> Maybe Int
m_add x y = do x' <- x -- Extract value from x
y' <- y -- Extract value from y
pure (x' + y') -- Add them
```

This function will extract the values from `x`

and `y`

, if they
are both available, or return `Nothing`

if one or both are not (“fail fast”). Managing the
`Nothing`

cases is achieved by the `>>=`

operator, hidden by the
`do`

notation.

```
*Interfaces> m_add (Just 20) (Just 22)
Just 42 : Maybe Int
*Interfaces> m_add (Just 20) Nothing
Nothing : Maybe Int
```

##### Pattern Matching Bind¶

Sometimes we want to pattern match immediately on the result of a function
in `do`

notation. For example, let’s say we have a function `readNumber`

which reads a number from the console, returning a value of the form
`Just x`

if the number is valid, or `Nothing`

otherwise:

```
readNumber : IO (Maybe Nat)
readNumber = do
input <- getLine
if all isDigit (unpack input)
then pure (Just (cast input))
else pure Nothing
```

If we then use it to write a function to read two numbers, returning
`Nothing`

if neither are valid, then we would like to pattern match
on the result of `readNumber`

:

```
readNumbers : IO (Maybe (Nat, Nat))
readNumbers =
do x <- readNumber
case x of
Nothing => pure Nothing
Just x_ok => do y <- readNumber
case y of
Nothing => pure Nothing
Just y_ok => pure (Just (x_ok, y_ok))
```

If there’s a lot of error handling, this could get deeply nested very quickly!
So instead, we can combine the bind and the pattern match in one line. For example,
we could try pattern matching on values of the form `Just x_ok`

:

```
readNumbers : IO (Maybe (Nat, Nat))
readNumbers =
do Just x_ok <- readNumber
Just y_ok <- readNumber
pure (Just (x_ok, y_ok))
```

There is still a problem, however, because we’ve now omitted the case for
`Nothing`

so `readNumbers`

is no longer total! We can add the `Nothing`

case back as follows:

```
readNumbers : IO (Maybe (Nat, Nat))
readNumbers =
do Just x_ok <- readNumber | Nothing => pure Nothing
Just y_ok <- readNumber | Nothing => pure Nothing
pure (Just (x_ok, y_ok))
```

The effect of this version of `readNumbers`

is identical to the first (in
fact, it is syntactic sugar for it and directly translated back into that form).
The first part of each statement (`Just x_ok <-`

and `Just y_ok <-`

) gives
the preferred binding - if this matches, execution will continue with the rest
of the `do`

block. The second part gives the alternative bindings, of which
there may be more than one.

`!`

-notation¶

In many cases, using `do`

-notation can make programs unnecessarily
verbose, particularly in cases such as `m_add`

above where the value
bound is used once, immediately. In these cases, we can use a
shorthand version, as follows:

```
m_add : Maybe Int -> Maybe Int -> Maybe Int
m_add x y = pure (!x + !y)
```

The notation `!expr`

means that the expression `expr`

should be
evaluated and then implicitly bound. Conceptually, we can think of
`!`

as being a prefix function with the following type:

```
(!) : m a -> a
```

Note, however, that it is not really a function, merely syntax! In
practice, a subexpression `!expr`

will lift `expr`

as high as
possible within its current scope, bind it to a fresh name `x`

, and
replace `!expr`

with `x`

. Expressions are lifted depth first, left
to right. In practice, `!`

-notation allows us to program in a more
direct style, while still giving a notational clue as to which
expressions are monadic.

For example, the expression:

```
let y = 42 in f !(g !(print y) !x)
```

is lifted to:

```
let y = 42 in do y' <- print y
x' <- x
g' <- g y' x'
f g'
```

##### Monad comprehensions¶

The list comprehension notation we saw in Section
More Expressions is more general, and applies to anything which
has an implementation of both `Monad`

and `Alternative`

:

```
interface Applicative f => Alternative (f : Type -> Type) where
empty : f a
(<|>) : f a -> f a -> f a
```

In general, a comprehension takes the form ```
[ exp | qual1, qual2, …,
qualn ]
```

where `quali`

can be one of:

- A generator
`x <- e`

- A
*guard*, which is an expression of type`Bool`

- A let binding
`let x = e`

To translate a comprehension `[exp | qual1, qual2, …, qualn]`

, first
any qualifier `qual`

which is a *guard* is translated to ```
guard
qual
```

, using the following function:

```
guard : Alternative f => Bool -> f ()
```

Then the comprehension is converted to `do`

notation:

```
do { qual1; qual2; ...; qualn; pure exp; }
```

Using monad comprehensions, an alternative definition for `m_add`

would be:

```
m_add : Maybe Int -> Maybe Int -> Maybe Int
m_add x y = [ x' + y' | x' <- x, y' <- y ]
```

#### Idiom brackets¶

While `do`

notation gives an alternative meaning to sequencing,
idioms give an alternative meaning to *application*. The notation and
larger example in this section is inspired by Conor McBride and Ross
Paterson’s paper “Applicative Programming with Effects” [1].

First, let us revisit `m_add`

above. All it is really doing is
applying an operator to two values extracted from `Maybe Int`

. We
could abstract out the application:

```
m_app : Maybe (a -> b) -> Maybe a -> Maybe b
m_app (Just f) (Just a) = Just (f a)
m_app _ _ = Nothing
```

Using this, we can write an alternative `m_add`

which uses this
alternative notion of function application, with explicit calls to
`m_app`

:

```
m_add' : Maybe Int -> Maybe Int -> Maybe Int
m_add' x y = m_app (m_app (Just (+)) x) y
```

Rather than having to insert `m_app`

everywhere there is an
application, we can use idiom brackets to do the job for us.
To do this, we can give `Maybe`

an implementation of `Applicative`

as follows, where `<*>`

is defined in the same way as `m_app`

above (this is defined in the Idris library):

```
Applicative Maybe where
pure = Just
(Just f) <*> (Just a) = Just (f a)
_ <*> _ = Nothing
```

Using `<*>`

we can use this implementation as follows, where a function
application `[| f a1 …an |]`

is translated into ```
pure f <*> a1 <*>
… <*> an
```

:

```
m_add' : Maybe Int -> Maybe Int -> Maybe Int
m_add' x y = [| x + y |]
```

##### An error-handling interpreter¶

Idiom notation is commonly useful when defining evaluators. McBride and Paterson describe such an evaluator [1], for a language similar to the following:

```
data Expr = Var String -- variables
| Val Int -- values
| Add Expr Expr -- addition
```

Evaluation will take place relative to a context mapping variables
(represented as `String`

s) to `Int`

values, and can possibly fail.
We define a data type `Eval`

to wrap an evaluator:

```
data Eval : Type -> Type where
MkEval : (List (String, Int) -> Maybe a) -> Eval a
```

Wrapping the evaluator in a data type means we will be able to provide implementations of interfaces for it later. We begin by defining a function to retrieve values from the context during evaluation:

```
fetch : String -> Eval Int
fetch x = MkEval (\e => fetchVal e) where
fetchVal : List (String, Int) -> Maybe Int
fetchVal [] = Nothing
fetchVal ((v, val) :: xs) = if (x == v)
then (Just val)
else (fetchVal xs)
```

When defining an evaluator for the language, we will be applying functions in
the context of an `Eval`

, so it is natural to give `Eval`

an implementation
of `Applicative`

. Before `Eval`

can have an implementation of
`Applicative`

it is necessary for `Eval`

to have an implementation of
`Functor`

:

```
Functor Eval where
map f (MkEval g) = MkEval (\e => map f (g e))
Applicative Eval where
pure x = MkEval (\e => Just x)
(<*>) (MkEval f) (MkEval g) = MkEval (\x => app (f x) (g x)) where
app : Maybe (a -> b) -> Maybe a -> Maybe b
app (Just fx) (Just gx) = Just (fx gx)
app _ _ = Nothing
```

Evaluating an expression can now make use of the idiomatic application to handle errors:

```
eval : Expr -> Eval Int
eval (Var x) = fetch x
eval (Val x) = [| x |]
eval (Add x y) = [| eval x + eval y |]
runEval : List (String, Int) -> Expr -> Maybe Int
runEval env e = case eval e of
MkEval envFn => envFn env
```

#### Named Implementations¶

It can be desirable to have multiple implementations of an interface for the
same type, for example to provide alternative methods for sorting or printing
values. To achieve this, implementations can be *named* as follows:

```
[myord] Ord Nat where
compare Z (S n) = GT
compare (S n) Z = LT
compare Z Z = EQ
compare (S x) (S y) = compare @{myord} x y
```

This declares an implementation as normal, but with an explicit name,
`myord`

. The syntax `compare @{myord}`

gives an explicit implementation to
`compare`

, otherwise it would use the default implementation for `Nat`

. We
can use this, for example, to sort a list of `Nat`

in reverse.
Given the following list:

```
testList : List Nat
testList = [3,4,1]
```

We can sort it using the default `Ord`

implementation, then the named
implementation `myord`

as follows, at the Idris prompt:

```
*named_impl> show (sort testList)
"[sO, sssO, ssssO]" : String
*named_impl> show (sort @{myord} testList)
"[ssssO, sssO, sO]" : String
```

Sometimes, we also need access to a named parent implementation. For example,
the prelude defines the following `Semigroup`

interface:

```
interface Semigroup ty where
(<+>) : ty -> ty -> ty
```

Then it defines `Monoid`

, which extends `Semigroup`

with a “neutral”
value:

```
interface Semigroup ty => Monoid ty where
neutral : ty
```

We can define two different implementations of `Semigroup`

and
`Monoid`

for `Nat`

, one based on addition and one on multiplication:

```
[PlusNatSemi] Semigroup Nat where
(<+>) x y = x + y
[MultNatSemi] Semigroup Nat where
(<+>) x y = x * y
```

The neutral value for addition is `0`

, but the neutral value for multiplication
is `1`

. It’s important, therefore, that when we define implementations
of `Monoid`

they extend the correct `Semigroup`

implementation. We can
do this with a `using`

clause in the implementation as follows:

```
[PlusNatMonoid] Monoid Nat using PlusNatSemi where
neutral = 0
[MultNatMonoid] Monoid Nat using MultNatSemi where
neutral = 1
```

The `using PlusNatSemi`

clause indicates that `PlusNatMonoid`

should
extend `PlusNatSemi`

specifically.

#### Determining Parameters¶

When an interface has more than one parameter, it can help resolution if the parameters used to find an implementation are restricted. For example:

```
interface Monad m => MonadState s (m : Type -> Type) | m where
get : m s
put : s -> m ()
```

In this interface, only `m`

needs to be known to find an implementation of
this interface, and `s`

can then be determined from the implementation. This
is declared with the `| m`

after the interface declaration. We call `m`

a
*determining parameter* of the `MonadState`

interface, because it is the
parameter used to find an implementation.

[1] | (1, 2) Conor McBride and Ross Paterson. 2008. Applicative programming
with effects. J. Funct. Program. 18, 1 (January 2008),
1-13. DOI=10.1017/S0956796807006326
http://dx.doi.org/10.1017/S0956796807006326 |

### Modules and Namespaces¶

An Idris program consists of a collection of modules. Each module
includes an optional `module`

declaration giving the name of the
module, a list of `import`

statements giving the other modules which
are to be imported, and a collection of declarations and definitions of
types, interfaces and functions. For example, the listing below gives a
module which defines a binary tree type `BTree`

(in a file
`Btree.idr`

):

```
module Btree
public export
data BTree a = Leaf
| Node (BTree a) a (BTree a)
export
insert : Ord a => a -> BTree a -> BTree a
insert x Leaf = Node Leaf x Leaf
insert x (Node l v r) = if (x < v) then (Node (insert x l) v r)
else (Node l v (insert x r))
export
toList : BTree a -> List a
toList Leaf = []
toList (Node l v r) = Btree.toList l ++ (v :: Btree.toList r)
export
toTree : Ord a => List a -> BTree a
toTree [] = Leaf
toTree (x :: xs) = insert x (toTree xs)
```

The modifiers `export`

and `public export`

say which names are visible
from other modules. These are explained further below.

Then, this gives a main program (in a file
`bmain.idr`

) which uses the `Btree`

module to sort a list:

```
module Main
import Btree
main : IO ()
main = do let t = toTree [1,8,2,7,9,3]
print (Btree.toList t)
```

The same names can be defined in multiple modules: names are *qualified* with
the name of the module. The names defined in the `Btree`

module are, in full:

`Btree.BTree`

`Btree.Leaf`

`Btree.Node`

`Btree.insert`

`Btree.toList`

`Btree.toTree`

If names are otherwise unambiguous, there is no need to give the fully qualified name. Names can be disambiguated either by giving an explicit qualification, or according to their type.

There is no formal link between the module name and its filename,
although it is generally advisable to use the same name for each. An
`import`

statement refers to a filename, using dots to separate
directories. For example, `import foo.bar`

would import the file
`foo/bar.idr`

, which would conventionally have the module declaration
`module foo.bar`

. The only requirement for module names is that the
main module, with the `main`

function, must be called
`Main`

— although its filename need not be `Main.idr`

.

#### Export Modifiers¶

Idris allows for fine-grained control over the visibility of a
module’s contents. By default, all names defined in a module are kept
private. This aides in specification of a minimal interface and for
internal details to be left hidden. Idris allows for functions,
types, and interfaces to be marked as: `private`

, `export`

, or
`public export`

. Their general meaning is as follows:

`private`

meaning that it’s not exported at all. This is the default.`export`

meaning that its top level type is exported.`public export`

meaning that the entire definition is exported.

A further restriction in modifying the visibility is that definitions
must not refer to anything within a lower level of visibility. For
example, `public export`

definitions cannot use private names, and
`export`

types cannot use private names. This is to prevent private
names leaking into a module’s interface.

##### Meaning for Functions¶

`export`

the type is exported`public export`

the type and definition are exported, and the definition can be used after it is imported. In other words, the definition itself is considered part of the module’s interface. The long name`public export`

is intended to make you think twice about doing this.

Note

Type synonyms in Idris are created by writing a function. When
setting the visibility for a module, it might be a good idea to
`public export`

all type synonyms if they are to be used outside
the module. Otherwise, Idris won’t know what the synonym is a
synonym for.

Since `public export`

means that a function’s definition is exported,
this effectively makes the function definition part of the module’s API.
Therefore, it’s generally a good idea to avoid using `public export`

for
functions unless you really mean to export the full definition.

##### Meaning for Data Types¶

For data types, the meanings are:

`export`

the type constructor is exported`public export`

the type constructor and data constructors are exported

##### Meaning for Interfaces¶

For interfaces, the meanings are:

`export`

the interface name is exported`public export`

the interface name, method names and default definitions are exported

`%access`

Directive¶

The default export mode can be changed with the `%access`

directive, for example:

```
module Btree
%access export
public export
data BTree a = Leaf
| Node (BTree a) a (BTree a)
insert : Ord a => a -> BTree a -> BTree a
insert x Leaf = Node Leaf x Leaf
insert x (Node l v r) = if (x < v) then (Node (insert x l) v r)
else (Node l v (insert x r))
toList : BTree a -> List a
toList Leaf = []
toList (Node l v r) = Btree.toList l ++ (v :: Btree.toList r)
toTree : Ord a => List a -> BTree a
toTree [] = Leaf
toTree (x :: xs) = insert x (toTree xs)
```

In this case, any function with no access modifier will be exported as
`export`

, rather than left `private`

.

##### Propagating Inner Module API’s¶

Additionally, a module can re-export a module it has imported, by using
the `public`

modifier on an `import`

. For example:

```
module A
import B
import public C
```

The module `A`

will export the name `a`

, as well as any public or
abstract names in module `C`

, but will not re-export anything from
module `B`

.

#### Explicit Namespaces¶

Defining a module also defines a namespace implicitly. However,
namespaces can also be given *explicitly*. This is most useful if you
wish to overload names within the same module:

```
module Foo
namespace x
test : Int -> Int
test x = x * 2
namespace y
test : String -> String
test x = x ++ x
```

This (admittedly contrived) module defines two functions with fully
qualified names `Foo.x.test`

and `Foo.y.test`

, which can be
disambiguated by their types:

```
*Foo> test 3
6 : Int
*Foo> test "foo"
"foofoo" : String
```

#### Parameterised blocks¶

Groups of functions can be parameterised over a number of arguments
using a `parameters`

declaration, for example:

```
parameters (x : Nat, y : Nat)
addAll : Nat -> Nat
addAll z = x + y + z
```

The effect of a `parameters`

block is to add the declared parameters
to every function, type and data constructor within the
block. Specifically, adding the parameters to the front of the
argument list. Outside the block, the parameters must be given
explicitly. The `addAll`

function, when called from the REPL, will
thus have the following type signature.

```
*params> :t addAll
addAll : Nat -> Nat -> Nat -> Nat
```

and the following definition.

```
addAll : (x : Nat) -> (y : Nat) -> (z : Nat) -> Nat
addAll x y z = x + y + z
```

Parameters blocks can be nested, and can also include data declarations, in which case the parameters are added explicitly to all type and data constructors. They may also be dependent types with implicit arguments:

```
parameters (y : Nat, xs : Vect x a)
data Vects : Type -> Type where
MkVects : Vect y a -> Vects a
append : Vects a -> Vect (x + y) a
append (MkVects ys) = xs ++ ys
```

To use `Vects`

or `append`

outside the block, we must also give the
`xs`

and `y`

arguments. Here, we can use placeholders for the values
which can be inferred by the type checker:

```
*params> show (append _ _ (MkVects _ [1,2,3] [4,5,6]))
"[1, 2, 3, 4, 5, 6]" : String
```

### Packages¶

Idris includes a simple build system for building packages and executables from a named package description file. These files can be used with the Idris compiler to manage the development process .

#### Package Descriptions¶

A package description includes the following:

- A header, consisting of the keyword
`package`

followed by a package name. Package names can be any valid Idris identifier. The iPKG format also takes a quoted version that accepts any valid filename. - Fields describing package contents,
`<field> = <value>`

.

At least one field must be the modules field, where the value is a
comma separated list of modules. For example, given an idris package
`maths`

that has modules `Maths.idr`

, `Maths.NumOps.idr`

,
`Maths.BinOps.idr`

, and `Maths.HexOps.idr`

, the corresponding
package file would be:

```
package maths
modules = Maths
, Maths.NumOps
, Maths.BinOps
, Maths.HexOps
```

Other examples of package files can be found in the `libs`

directory
of the main Idris repository, and in third-party libraries.

#### Using Package files¶

Idris itself is aware about packages, and special commands are
available to help with, for example, building packages, installing
packages, and cleaning packages. For instance, given the `maths`

package from earlier we can use Idris as follows:

`idris --build maths.ipkg`

will build all modules in the package`idris --install maths.ipkg`

will install the package, making it accessible by other Idris libraries and programs.`idris --clean maths.ipkg`

will delete all intermediate code and executable files generated when building.

Once the maths package has been installed, the command line option
`--package maths`

makes it accessible (abbreviated to `-p maths`

).
For example:

```
idris -p maths Main.idr
```

#### Testing Idris Packages¶

The integrated build system includes a simple testing framework.
This framework collects functions listed in the `ipkg`

file under `tests`

.
All test functions must return `IO ()`

.

When you enter `idris --testpkg yourmodule.ipkg`

,
the build system creates a temporary file in a fresh environment on your machine
by listing the `tests`

functions under a single `main`

function.
It compiles this temporary file to an executable and then executes it.

The tests themselves are responsible for reporting their success or failure.
Test functions commonly use `putStrLn`

to report test results.
The test framework does not impose any standards for reporting and consequently
does not aggregate test results.

For example, lets take the following list of functions that are defined in a
module called `NumOps`

for a sample package `maths`

:

```
module Maths.NumOps
%access export -- to make functions under test visible
double : Num a => a -> a
double a = a + a
triple : Num a => a -> a
triple a = a + double a
```

A simple test module, with a qualified name of `Test.NumOps`

can be declared as:

```
module Test.NumOps
import Maths.NumOps
%access export -- to make the test functions visible
assertEq : Eq a => (given : a) -> (expected : a) -> IO ()
assertEq g e = if g == e
then putStrLn "Test Passed"
else putStrLn "Test Failed"
assertNotEq : Eq a => (given : a) -> (expected : a) -> IO ()
assertNotEq g e = if not (g == e)
then putStrLn "Test Passed"
else putStrLn "Test Failed"
testDouble : IO ()
testDouble = assertEq (double 2) 4
testTriple : IO ()
testTriple = assertNotEq (triple 2) 5
```

The functions `assertEq`

and `assertNotEq`

are used to run expected passing,
and failing, equality tests. The actual tests are `testDouble`

and `testTriple`

,
and are declared in the `maths.ipkg`

file as follows:

```
package maths
modules = Maths.NumOps
, Test.NumOps
tests = Test.NumOps.testDouble
, Test.NumOps.testTriple
```

The testing framework can then be invoked using `idris --testpkg maths.ipkg`

:

```
> idris --testpkg maths.ipkg
Type checking ./Maths/NumOps.idr
Type checking ./Test/NumOps.idr
Type checking /var/folders/63/np5g0d5j54x1s0z12rf41wxm0000gp/T/idristests144128232716531729.idr
Test Passed
Test Passed
```

Note how both tests have reported success by printing `Test Passed`

as we arranged for with the `assertEq`

and `assertNoEq`

functions.

#### Package Dependencies Using Atom¶

If you are using the Atom editor and have a dependency on another package,
corresponding to for instance `import Lightyear`

or `import Pruviloj`

,
you need to let Atom know that it should be loaded. The easiest way to
accomplish that is with a .ipkg file. The general contents of an ipkg file
will be described in the next section of the tutorial, but for now here is
a simple recipe for this trivial case:

- Create a folder myProject.
- Add a file myProject.ipkg containing just a couple of lines:

```
package myProject
pkgs = pruviloj, lightyear
```

- In Atom, use the File menu to Open Folder myProject.

### Example: The Well-Typed Interpreter¶

In this section, we’ll use the features we’ve seen so far to write a
larger example, an interpreter for a simple functional programming
language, with variables, function application, binary operators and
an `if...then...else`

construct. We will use the dependent type
system to ensure that any programs which can be represented are
well-typed.

#### Representing Languages¶

First, let us define the types in the language. We have integers,
booleans, and functions, represented by `Ty`

:

```
data Ty = TyInt | TyBool | TyFun Ty Ty
```

We can write a function to translate these representations to a concrete Idris type — remember that types are first class, so can be calculated just like any other value:

```
interpTy : Ty -> Type
interpTy TyInt = Integer
interpTy TyBool = Bool
interpTy (TyFun A T) = interpTy A -> interpTy T
```

We’re going to define a representation of our language in such a way
that only well-typed programs can be represented. We’ll index the
representations of expressions by their type, **and** the types of
local variables (the context). The context can be represented using
the `Vect`

data type, and as it will be used regularly it will be
represented as an implicit argument. To do so we define everything in
a `using`

block (keep in mind that everything after this point needs
to be indented so as to be inside the `using`

block):

```
using (G:Vect n Ty)
```

Expressions are indexed by the types of the local variables, and the type of the expression itself:

```
data Expr : Vect n Ty -> Ty -> Type
```

The full representation of expressions is:

```
data HasType : (i : Fin n) -> Vect n Ty -> Ty -> Type where
Stop : HasType FZ (t :: G) t
Pop : HasType k G t -> HasType (FS k) (u :: G) t
data Expr : Vect n Ty -> Ty -> Type where
Var : HasType i G t -> Expr G t
Val : (x : Integer) -> Expr G TyInt
Lam : Expr (a :: G) t -> Expr G (TyFun a t)
App : Expr G (TyFun a t) -> Expr G a -> Expr G t
Op : (interpTy a -> interpTy b -> interpTy c) ->
Expr G a -> Expr G b -> Expr G c
If : Expr G TyBool ->
Lazy (Expr G a) ->
Lazy (Expr G a) -> Expr G a
```

The code above makes use of the `Vect`

and `Fin`

types from the
Idris standard library. We import them because they are not provided
in the prelude:

```
import Data.Vect
import Data.Fin
```

Since expressions are indexed by their type, we can read the typing rules of the language from the definitions of the constructors. Let us look at each constructor in turn.

We use a nameless representation for variables — they are *de Bruijn
indexed*. Variables are represented by a proof of their membership in
the context, `HasType i G T`

, which is a proof that variable `i`

in context `G`

has type `T`

. This is defined as follows:

```
data HasType : (i : Fin n) -> Vect n Ty -> Ty -> Type where
Stop : HasType FZ (t :: G) t
Pop : HasType k G t -> HasType (FS k) (u :: G) t
```

We can treat *Stop* as a proof that the most recently defined variable
is well-typed, and *Pop n* as a proof that, if the `n`

th most
recently defined variable is well-typed, so is the `n+1`

th. In
practice, this means we use `Stop`

to refer to the most recently
defined variable, `Pop Stop`

to refer to the next, and so on, via
the `Var`

constructor:

```
Var : HasType i G t -> Expr G t
```

So, in an expression `\x. \y. x y`

, the variable `x`

would have a
de Bruijn index of 1, represented as `Pop Stop`

, and `y 0`

,
represented as `Stop`

. We find these by counting the number of
lambdas between the definition and the use.

A value carries a concrete representation of an integer:

```
Val : (x : Integer) -> Expr G TyInt
```

A lambda creates a function. In the scope of a function of type ```
a ->
t
```

, there is a new local variable of type `a`

, which is expressed
by the context index:

```
Lam : Expr (a :: G) t -> Expr G (TyFun a t)
```

Function application produces a value of type `t`

given a function
from `a`

to `t`

and a value of type `a`

:

```
App : Expr G (TyFun a t) -> Expr G a -> Expr G t
```

We allow arbitrary binary operators, where the type of the operator informs what the types of the arguments must be:

```
Op : (interpTy a -> interpTy b -> interpTy c) ->
Expr G a -> Expr G b -> Expr G c
```

Finally, `If`

expressions make a choice given a boolean. Each branch
must have the same type, and we will evaluate the branches lazily so
that only the branch which is taken need be evaluated:

```
If : Expr G TyBool ->
Lazy (Expr G a) ->
Lazy (Expr G a) ->
Expr G a
```

#### Writing the Interpreter¶

When we evaluate an `Expr`

, we’ll need to know the values in scope,
as well as their types. `Env`

is an environment, indexed over the
types in scope. Since an environment is just another form of list,
albeit with a strongly specified connection to the vector of local
variable types, we use the usual `::`

and `Nil`

constructors so
that we can use the usual list syntax. Given a proof that a variable
is defined in the context, we can then produce a value from the
environment:

```
data Env : Vect n Ty -> Type where
Nil : Env Nil
(::) : interpTy a -> Env G -> Env (a :: G)
lookup : HasType i G t -> Env G -> interpTy t
lookup Stop (x :: xs) = x
lookup (Pop k) (x :: xs) = lookup k xs
```

Given this, an interpreter is a function which
translates an `Expr`

into a concrete Idris value with respect to a
specific environment:

```
interp : Env G -> Expr G t -> interpTy t
```

The complete interpreter is defined as follows, for reference. For each constructor, we translate it into the corresponding Idris value:

```
interp env (Var i) = lookup i env
interp env (Val x) = x
interp env (Lam sc) = \x => interp (x :: env) sc
interp env (App f s) = interp env f (interp env s)
interp env (Op op x y) = op (interp env x) (interp env y)
interp env (If x t e) = if interp env x then interp env t
else interp env e
```

Let us look at each case in turn. To translate a variable, we simply look it up in the environment:

```
interp env (Var i) = lookup i env
```

To translate a value, we just return the concrete representation of the value:

```
interp env (Val x) = x
```

Lambdas are more interesting. In this case, we construct a function which interprets the scope of the lambda with a new value in the environment. So, a function in the object language is translated to an Idris function:

```
interp env (Lam sc) = \x => interp (x :: env) sc
```

For an application, we interpret the function and its argument and apply
it directly. We know that interpreting `f`

must produce a function,
because of its type:

```
interp env (App f s) = interp env f (interp env s)
```

Operators and conditionals are, again, direct translations into the
equivalent Idris constructs. For operators, we apply the function to
its operands directly, and for `If`

, we apply the Idris
`if...then...else`

construct directly.

```
interp env (Op op x y) = op (interp env x) (interp env y)
interp env (If x t e) = if interp env x then interp env t
else interp env e
```

#### Testing¶

We can make some simple test functions. Firstly, adding two inputs
`\x. \y. y + x`

is written as follows:

```
add : Expr G (TyFun TyInt (TyFun TyInt TyInt))
add = Lam (Lam (Op (+) (Var Stop) (Var (Pop Stop))))
```

More interestingly, a factorial function `fact`

(e.g. `\x. if (x == 0) then 1 else (fact (x-1) * x)`

),
can be written as:

```
fact : Expr G (TyFun TyInt TyInt)
fact = Lam (If (Op (==) (Var Stop) (Val 0))
(Val 1)
(Op (*) (App fact (Op (-) (Var Stop) (Val 1)))
(Var Stop)))
```

#### Running¶

To finish, we write a `main`

program which interprets the factorial
function on user input:

```
main : IO ()
main = do putStr "Enter a number: "
x <- getLine
printLn (interp [] fact (cast x))
```

Here, `cast`

is an overloaded function which converts a value from
one type to another if possible. Here, it converts a string to an
integer, giving 0 if the input is invalid. An example run of this
program at the Idris interactive environment is:

```
$ idris interp.idr
____ __ _
/ _/___/ /____(_)____
/ // __ / ___/ / ___/ Version 1.3.0
_/ // /_/ / / / (__ ) http://www.idris-lang.org/
/___/\__,_/_/ /_/____/ Type :? for help
Type checking ./interp.idr
*interp> :exec
Enter a number: 6
720
*interp>
```

##### Aside: `cast`

¶

The prelude defines an interface `Cast`

which allows conversion
between types:

```
interface Cast from to where
cast : from -> to
```

It is a *multi-parameter* interface, defining the source type and
object type of the cast. It must be possible for the type checker to
infer *both* parameters at the point where the cast is applied. There
are casts defined between all of the primitive types, as far as they
make sense.

### Views and the “`with`

” rule¶

#### Dependent pattern matching¶

Since types can depend on values, the form of some arguments can be
determined by the value of others. For example, if we were to write
down the implicit length arguments to `(++)`

, we’d see that the form
of the length argument was determined by whether the vector was empty
or not:

```
(++) : Vect n a -> Vect m a -> Vect (n + m) a
(++) {n=Z} [] ys = ys
(++) {n=S k} (x :: xs) ys = x :: xs ++ ys
```

If `n`

was a successor in the `[]`

case, or zero in the `::`

case, the definition would not be well typed.

#### The `with`

rule — matching intermediate values¶

Very often, we need to match on the result of an intermediate
computation. Idris provides a construct for this, the `with`

rule, inspired by views in `Epigram`

[1], which takes account of
the fact that matching on a value in a dependently typed language can
affect what we know about the forms of other values. In its simplest
form, the `with`

rule adds another argument to the function being
defined.

We have already seen a vector filter function. This time, we define it
using `with`

as follows:

```
filter : (a -> Bool) -> Vect n a -> (p ** Vect p a)
filter p [] = ( _ ** [] )
filter p (x :: xs) with (filter p xs)
filter p (x :: xs) | ( _ ** xs' ) = if (p x) then ( _ ** x :: xs' ) else ( _ ** xs' )
```

Here, the `with`

clause allows us to deconstruct the result of
`filter p xs`

. The view refined argument pattern ```
filter p (x ::
xs)
```

goes beneath the `with`

clause, followed by a vertical bar
`|`

, followed by the deconstructed intermediate result ```
( _ ** xs'
)
```

. If the view refined argument pattern is unchanged from the
original function argument pattern, then the left side of `|`

is
extraneous and may be omitted:

```
filter p (x :: xs) with (filter p xs)
| ( _ ** xs' ) = if (p x) then ( _ ** x :: xs' ) else ( _ ** xs' )
```

If the intermediate computation itself has a dependent type, then the
result can affect the forms of other arguments — we can learn the form
of one value by testing another. In these cases, view refined argument
patterns must be explicit. For example, a `Nat`

is either even or
odd. If it is even it will be the sum of two equal `Nat`

.
Otherwise, it is the sum of two equal `Nat`

plus one:

```
data Parity : Nat -> Type where
Even : Parity (n + n)
Odd : Parity (S (n + n))
```

We say `Parity`

is a *view* of `Nat`

. It has a *covering function*
which tests whether it is even or odd and constructs the predicate
accordingly.

```
parity : (n:Nat) -> Parity n
```

We’ll come back to the definition of `parity`

shortly. We can use it
to write a function which converts a natural number to a list of
binary digits (least significant first) as follows, using the `with`

rule:

```
natToBin : Nat -> List Bool
natToBin Z = Nil
natToBin k with (parity k)
natToBin (j + j) | Even = False :: natToBin j
natToBin (S (j + j)) | Odd = True :: natToBin j
```

The value of `parity k`

affects the form of `k`

, because the
result of `parity k`

depends on `k`

. So, as well as the patterns
for the result of the intermediate computation (`Even`

and `Odd`

)
right of the `|`

, we also write how the results affect the other
patterns left of the `|`

. That is:

- When
`parity k`

evaluates to`Even`

, we can refine the original argument`k`

to a refined pattern`(j + j)`

according to`Parity (n + n)`

from the`Even`

constructor definition. So`(j + j)`

replaces`k`

on the left side of`|`

, and the`Even`

constructor appears on the right side. The natural number`j`

in the refined pattern can be used on the right side of the`=`

sign. - Otherwise, when
`parity k`

evaluates to`Odd`

, the original argument`k`

is refined to`S (j + j)`

according to`Parity (S (n + n))`

from the`Odd`

constructor definition, and`Odd`

now appears on the right side of`|`

, again with the natural number`j`

used on the right side of the`=`

sign.

Note that there is a function in the patterns (`+`

) and repeated
occurrences of `j`

- this is allowed because another argument has
determined the form of these patterns.

We will return to this function in the next section Theorems in Practice to
complete the definition of `parity`

.

#### With and proofs¶

To use a dependent pattern match for theorem proving, it is sometimes necessary
to explicitly construct the proof resulting from the pattern match.
To do this, you can postfix the with clause with `proof p`

and the proof
generated by the pattern match will be in scope and named `p`

. For example:

```
data Foo = FInt Int | FBool Bool
optional : Foo -> Maybe Int
optional (FInt x) = Just x
optional (FBool b) = Nothing
isFInt : (foo:Foo) -> Maybe (x : Int ** (optional foo = Just x))
isFInt foo with (optional foo) proof p
isFInt foo | Nothing = Nothing -- here, p : Nothing = optional foo
isFInt foo | (Just x) = Just (x ** Refl) -- here, p : Just x = optional foo
```

[1] | Conor McBride and James McKinna. 2004. The view from the left. J. Funct. Program. 14, 1 (January 2004), 69-111. DOI=10.1017/S0956796803004829 http://dx.doi.org/10.1017/S0956796803004829ñ |

### Theorem Proving¶

#### Equality¶

Idris allows propositional equalities to be declared, allowing theorems about programs to be stated and proved. Equality is built in, but conceptually has the following definition:

```
data (=) : a -> b -> Type where
Refl : x = x
```

Equalities can be proposed between any values of any types, but the only way to construct a proof of equality is if values actually are equal. For example:

```
fiveIsFive : 5 = 5
fiveIsFive = Refl
twoPlusTwo : 2 + 2 = 4
twoPlusTwo = Refl
```

#### The Empty Type¶

There is an empty type, , which has no constructors. It is therefore impossible to construct an element of the empty type, at least without using a partially defined or general recursive function (see Section Totality Checking for more details). We can therefore use the empty type to prove that something is impossible, for example zero is never equal to a successor:

```
disjoint : (n : Nat) -> Z = S n -> Void
disjoint n p = replace {P = disjointTy} p ()
where
disjointTy : Nat -> Type
disjointTy Z = ()
disjointTy (S k) = Void
```

There is no need to worry too much about how this function works —
essentially, it applies the library function `replace`

, which uses an
equality proof to transform a predicate. Here we use it to transform a
value of a type which can exist, the empty tuple, to a value of a type
which can’t, by using a proof of something which can’t exist.

Once we have an element of the empty type, we can prove anything.
`void`

is defined in the library, to assist with proofs by
contradiction.

```
void : Void -> a
```

#### Simple Theorems¶

When type checking dependent types, the type itself gets *normalised*.
So imagine we want to prove the following theorem about the reduction
behaviour of `plus`

:

```
plusReduces : (n:Nat) -> plus Z n = n
```

We’ve written down the statement of the theorem as a type, in just the same way as we would write the type of a program. In fact there is no real distinction between proofs and programs. A proof, as far as we are concerned here, is merely a program with a precise enough type to guarantee a particular property of interest.

We won’t go into details here, but the Curry-Howard correspondence [1]
explains this relationship. The proof itself is trivial, because
`plus Z n`

normalises to `n`

by the definition of `plus`

:

```
plusReduces n = Refl
```

It is slightly harder if we try the arguments the other way, because
plus is defined by recursion on its first argument. The proof also works
by recursion on the first argument to `plus`

, namely `n`

.

```
plusReducesZ : (n:Nat) -> n = plus n Z
plusReducesZ Z = Refl
plusReducesZ (S k) = cong (plusReducesZ k)
```

`cong`

is a function defined in the library which states that equality
respects function application:

```
cong : {f : t -> u} -> a = b -> f a = f b
```

We can do the same for the reduction behaviour of plus on successors:

```
plusReducesS : (n:Nat) -> (m:Nat) -> S (plus n m) = plus n (S m)
plusReducesS Z m = Refl
plusReducesS (S k) m = cong (plusReducesS k m)
```

Even for trivial theorems like these, the proofs are a little tricky to construct in one go. When things get even slightly more complicated, it becomes too much to think about to construct proofs in this “batch mode”.

Idris provides interactive editing capabilities, which can help with building proofs. For more details on building proofs interactively in an editor, see Theorem Proving.

#### Theorems in Practice¶

The need to prove theorems can arise naturally in practice. For example,
previously (Views and the “with” rule) we implemented `natToBin`

using a function
`parity`

:

```
parity : (n:Nat) -> Parity n
```

However, we didn’t provide a definition for `parity`

. We might expect it
to look something like the following:

```
parity : (n:Nat) -> Parity n
parity Z = Even {n=Z}
parity (S Z) = Odd {n=Z}
parity (S (S k)) with (parity k)
parity (S (S (j + j))) | Even = Even {n=S j}
parity (S (S (S (j + j)))) | Odd = Odd {n=S j}
```

Unfortunately, this fails with a type error:

```
When checking right hand side of with block in views.parity with expected type
Parity (S (S (j + j)))
Type mismatch between
Parity (S j + S j) (Type of Even)
and
Parity (S (S (plus j j))) (Expected type)
```

The problem is that normalising `S j + S j`

, in the type of `Even`

doesn’t result in what we need for the type of the right hand side of
`Parity`

. We know that `S (S (plus j j))`

is going to be equal to
`S j + S j`

, but we need to explain it to Idris with a proof. We can
begin by adding some *holes* (see Holes) to the definition:

```
parity : (n:Nat) -> Parity n
parity Z = Even {n=Z}
parity (S Z) = Odd {n=Z}
parity (S (S k)) with (parity k)
parity (S (S (j + j))) | Even = let result = Even {n=S j} in
?helpEven
parity (S (S (S (j + j)))) | Odd = let result = Odd {n=S j} in
?helpOdd
```

Checking the type of `helpEven`

shows us what we need to prove for the
`Even`

case:

```
j : Nat
result : Parity (S (plus j (S j)))
--------------------------------------
helpEven : Parity (S (S (plus j j)))
```

We can therefore write a helper function to *rewrite* the type to the form
we need:

```
helpEven : (j : Nat) -> Parity (S j + S j) -> Parity (S (S (plus j j)))
helpEven j p = rewrite plusSuccRightSucc j j in p
```

The `rewrite ... in`

syntax allows you to change the required type of an
expression by rewriting it according to an equality proof. Here, we have
used `plusSuccRightSucc`

, which has the following type:

```
plusSuccRightSucc : (left : Nat) -> (right : Nat) -> S (left + right) = left + S right
```

We can see the effect of `rewrite`

by replacing the right hand side of
`helpEven`

with a hole, and working step by step. Beginning with the following:

```
helpEven : (j : Nat) -> Parity (S j + S j) -> Parity (S (S (plus j j)))
helpEven j p = ?helpEven_rhs
```

We can look at the type of `helpEven_rhs`

:

```
j : Nat
p : Parity (S (plus j (S j)))
--------------------------------------
helpEven_rhs : Parity (S (S (plus j j)))
```

Then we can `rewrite`

by applying `plusSuccRightSucc j j`

, which gives
an equation `S (j + j) = j + S j`

, thus replacing `S (j + j)`

(or,
in this case, `S (plus j j)`

since `S (j + j)`

reduces to that) in the
type with `j + S j`

:

```
helpEven : (j : Nat) -> Parity (S j + S j) -> Parity (S (S (plus j j)))
helpEven j p = rewrite plusSuccRightSucc j j in ?helpEven_rhs
```

Checking the type of `helpEven_rhs`

now shows what has happened, including
the type of the equation we just used (as the type of `_rewrite_rule`

):

```
j : Nat
p : Parity (S (plus j (S j)))
_rewrite_rule : S (plus j j) = plus j (S j)
--------------------------------------
helpEven_rhs : Parity (S (plus j (S j)))
```

Using `rewrite`

and another helper for the `Odd`

case, we can complete
`parity`

as follows:

```
helpEven : (j : Nat) -> Parity (S j + S j) -> Parity (S (S (plus j j)))
helpEven j p = rewrite plusSuccRightSucc j j in p
helpOdd : (j : Nat) -> Parity (S (S (j + S j))) -> Parity (S (S (S (j + j))))
helpOdd j p = rewrite plusSuccRightSucc j j in p
parity : (n:Nat) -> Parity n
parity Z = Even {n=Z}
parity (S Z) = Odd {n=Z}
parity (S (S k)) with (parity k)
parity (S (S (j + j))) | Even = helpEven j (Even {n = S j})
parity (S (S (S (j + j)))) | Odd = helpOdd j (Odd {n = S j})
```

Full details of `rewrite`

are beyond the scope of this introductory tutorial,
but it is covered in the theorem proving tutorial (see Theorem Proving).

#### Totality Checking¶

If we really want to trust our proofs, it is important that they are
defined by *total* functions — that is, a function which is defined for
all possible inputs and is guaranteed to terminate. Otherwise we could
construct an element of the empty type, from which we could prove
anything:

```
-- making use of 'hd' being partially defined
empty1 : Void
empty1 = hd [] where
hd : List a -> a
hd (x :: xs) = x
-- not terminating
empty2 : Void
empty2 = empty2
```

Internally, Idris checks every definition for totality, and we can check at
the prompt with the `:total`

command. We see that neither of the above
definitions is total:

```
*Theorems> :total empty1
possibly not total due to: empty1#hd
not total as there are missing cases
*Theorems> :total empty2
possibly not total due to recursive path empty2
```

Note the use of the word “possibly” — a totality check can, of course, never be certain due to the undecidability of the halting problem. The check is, therefore, conservative. It is also possible (and indeed advisable, in the case of proofs) to mark functions as total so that it will be a compile time error for the totality check to fail:

```
total empty2 : Void
empty2 = empty2
```

```
Type checking ./theorems.idr
theorems.idr:25:empty2 is possibly not total due to recursive path empty2
```

Reassuringly, our proof in Section The Empty Type that the zero and successor constructors are disjoint is total:

```
*theorems> :total disjoint
Total
```

The totality check is, necessarily, conservative. To be recorded as
total, a function `f`

must:

- Cover all possible inputs
- Be
*well-founded*— i.e. by the time a sequence of (possibly mutually) recursive calls reaches`f`

again, it must be possible to show that one of its arguments has decreased. - Not use any data types which are not
*strictly positive* - Not call any non-total functions

##### Directives and Compiler Flags for Totality¶

By default, Idris allows all well-typed definitions, whether total or not. However, it is desirable for functions to be total as far as possible, as this provides a guarantee that they provide a result for all possible inputs, in finite time. It is possible to make total functions a requirement, either:

- By using the
`--total`

compiler flag. - By adding a
`%default total`

directive to a source file. All definitions after this will be required to be total, unless explicitly flagged as`partial`

.

All functions *after* a `%default total`

declaration are required to
be total. Correspondingly, after a `%default partial`

declaration, the
requirement is relaxed.

Finally, the compiler flag `--warnpartial`

causes to print a warning
for any undeclared partial function.

##### Totality checking issues¶

Please note that the totality checker is not perfect! Firstly, it is
necessarily conservative due to the undecidability of the halting
problem, so many programs which *are* total will not be detected as
such. Secondly, the current implementation has had limited effort put
into it so far, so there may still be cases where it believes a function
is total which is not. Do not rely on it for your proofs yet!

##### Hints for totality¶

In cases where you believe a program is total, but Idris does not agree, it is
possible to give hints to the checker to give more detail for a termination
argument. The checker works by ensuring that all chains of recursive calls
eventually lead to one of the arguments decreasing towards a base case, but
sometimes this is hard to spot. For example, the following definition cannot be
checked as `total`

because the checker cannot decide that `filter (< x) xs`

will always be smaller than `(x :: xs)`

:

```
qsort : Ord a => List a -> List a
qsort [] = []
qsort (x :: xs)
= qsort (filter (< x) xs) ++
(x :: qsort (filter (>= x) xs))
```

The function `assert_smaller`

, defined in the prelude, is intended to
address this problem:

```
assert_smaller : a -> a -> a
assert_smaller x y = y
```

It simply evaluates to its second argument, but also asserts to the
totality checker that `y`

is structurally smaller than `x`

. This can
be used to explain the reasoning for totality if the checker cannot work
it out itself. The above example can now be written as:

```
total
qsort : Ord a => List a -> List a
qsort [] = []
qsort (x :: xs)
= qsort (assert_smaller (x :: xs) (filter (< x) xs)) ++
(x :: qsort (assert_smaller (x :: xs) (filter (>= x) xs)))
```

The expression `assert_smaller (x :: xs) (filter (<= x) xs)`

asserts
that the result of the filter will always be smaller than the pattern
`(x :: xs)`

.

In more extreme cases, the function `assert_total`

marks a
subexpression as always being total:

```
assert_total : a -> a
assert_total x = x
```

In general, this function should be avoided, but it can be very useful when reasoning about primitives or externally defined functions (for example from a C library) where totality can be shown by an external argument.

[1] | Timothy G. Griffin. 1989. A formulae-as-type notion of control. In Proceedings of the 17th ACM SIGPLAN-SIGACT symposium on Principles of programming languages (POPL ‘90). ACM, New York, NY, USA, 47-58. DOI=10.1145/96709.96714 http://doi.acm.org/10.1145/96709.96714 |

### Provisional Definitions¶

Sometimes when programming with dependent types, the type required by
the type checker and the type of the program we have written will be
different (in that they do not have the same normal form), but
nevertheless provably equal. For example, recall the `parity`

function:

```
data Parity : Nat -> Type where
Even : Parity (n + n)
Odd : Parity (S (n + n))
```

We’d like to implement this as follows:

```
parity : (n:Nat) -> Parity n
parity Z = Even {n=Z}
parity (S Z) = Odd {n=Z}
parity (S (S k)) with (parity k)
parity (S (S (j + j))) | Even = Even {n=S j}
parity (S (S (S (j + j)))) | Odd = Odd {n=S j}
```

This simply states that zero is even, one is odd, and recursively, the
parity of `k+2`

is the same as the parity of `k`

. Explicitly marking
the value of `n`

is even and odd is necessary to help type inference.
Unfortunately, the type checker rejects this:

```
viewsbroken.idr:12:10:When elaborating right hand side of ViewsBroken.parity:
Type mismatch between
Parity (plus (S j) (S j))
and
Parity (S (S (plus j j)))
Specifically:
Type mismatch between
plus (S j) (S j)
and
S (S (plus j j))
```

The type checker is telling us that `(j+1)+(j+1)`

and `2+j+j`

do not
normalise to the same value. This is because `plus`

is defined by
recursion on its first argument, and in the second value, there is a
successor symbol on the second argument, so this will not help with
reduction. These values are obviously equal — how can we rewrite the
program to fix this problem?

#### Provisional definitions¶

*Provisional definitions* help with this problem by allowing us to defer
the proof details until a later point. There are two main reasons why
they are useful.

- When
*prototyping*, it is useful to be able to test programs before finishing all the details of proofs. - When
*reading*a program, it is often much clearer to defer the proof details so that they do not distract the reader from the underlying algorithm.

Provisional definitions are written in the same way as ordinary
definitions, except that they introduce the right hand side with a
`?=`

rather than `=`

. We define `parity`

as follows:

```
parity : (n:Nat) -> Parity n
parity Z = Even {n=Z}
parity (S Z) = Odd {n=Z}
parity (S (S k)) with (parity k)
parity (S (S (j + j))) | Even ?= Even {n=S j}
parity (S (S (S (j + j)))) | Odd ?= Odd {n=S j}
```

When written in this form, instead of reporting a type error, Idris will insert a hole standing for a theorem which will correct the type error. Idris tells us we have two proof obligations, with names generated from the module and function names:

```
*views> :m
Global holes:
[views.parity_lemma_2,views.parity_lemma_1]
```

The first of these has the following type:

```
*views> :p views.parity_lemma_1
---------------------------------- (views.parity_lemma_1) --------
{hole0} : (j : Nat) -> (Parity (plus (S j) (S j))) -> Parity (S (S (plus j j)))
-views.parity_lemma_1>
```

The two arguments are `j`

, the variable in scope from the pattern
match, and `value`

, which is the value we gave in the right hand side
of the provisional definition. Our goal is to rewrite the type so that
we can use this value. We can achieve this using the following theorem
from the prelude:

```
plusSuccRightSucc : (left : Nat) -> (right : Nat) ->
S (left + right) = left + (S right)
```

We need to use `compute`

again to unfold the definition of `plus`

:

```
-views.parity_lemma_1> compute
---------------------------------- (views.parity_lemma_1) --------
{hole0} : (j : Nat) -> (Parity (S (plus j (S j)))) -> Parity (S (S (plus j j)))
```

After applying `intros`

we have:

```
-views.parity_lemma_1> intros
j : Nat
value : Parity (S (plus j (S j)))
---------------------------------- (views.parity_lemma_1) --------
{hole2} : Parity (S (S (plus j j)))
```

Then we apply the `plusSuccRightSucc`

rewrite rule, symmetrically, to
`j`

and `j`

, giving:

```
-views.parity_lemma_1> rewrite sym (plusSuccRightSucc j j)
j : Nat
value : Parity (S (plus j (S j)))
---------------------------------- (views.parity_lemma_1) --------
{hole3} : Parity (S (plus j (S j)))
```

`sym`

is a function, defined in the library, which reverses the order
of the rewrite:

```
sym : l = r -> r = l
sym Refl = Refl
```

We can complete this proof using the `trivial`

tactic, which finds
`value`

in the premises. The proof of the second lemma proceeds in
exactly the same way.

We can now test the `natToBin`

function from Section The with rule — matching intermediate values
at the prompt. The number 42 is 101010 in binary. The binary digits are
reversed:

```
*views> show (natToBin 42)
"[False, True, False, True, False, True]" : String
```

#### Suspension of Disbelief¶

Idris requires that proofs be complete before compiling programs (although evaluation at the prompt is possible without proof details). Sometimes, especially when prototyping, it is easier not to have to do this. It might even be beneficial to test programs before attempting to prove things about them — if testing finds an error, you know you had better not waste your time proving something!

Therefore, Idris provides a built-in coercion function, which allows you to use a value of the incorrect types:

```
believe_me : a -> b
```

Obviously, this should be used with extreme caution. It is useful when
prototyping, and can also be appropriate when asserting properties of
external code (perhaps in an external C library). The “proof” of
`views.parity_lemma_1`

using this is:

```
views.parity_lemma_2 = proof {
intro;
intro;
exact believe_me value;
}
```

The `exact`

tactic allows us to provide an exact value for the proof.
In this case, we assert that the value we gave was correct.

#### Example: Binary numbers¶

Previously, we implemented conversion to binary numbers using the
`Parity`

view. Here, we show how to use the same view to implement a
verified conversion to binary. We begin by indexing binary numbers over
their `Nat`

equivalent. This is a common pattern, linking a
representation (in this case `Binary`

) with a meaning (in this case
`Nat`

):

```
data Binary : Nat -> Type where
BEnd : Binary Z
BO : Binary n -> Binary (n + n)
BI : Binary n -> Binary (S (n + n))
```

`BO`

and `BI`

take a binary number as an argument and effectively
shift it one bit left, adding either a zero or one as the new least
significant bit. The index, `n + n`

or `S (n + n)`

states the result
that this left shift then add will have to the meaning of the number.
This will result in a representation with the least significant bit at
the front.

Now a function which converts a Nat to binary will state, in the type, that the resulting binary number is a faithful representation of the original Nat:

```
natToBin : (n:Nat) -> Binary n
```

The `Parity`

view makes the definition fairly simple — halving the
number is effectively a right shift after all — although we need to use
a provisional definition in the Odd case:

```
natToBin : (n:Nat) -> Binary n
natToBin Z = BEnd
natToBin (S k) with (parity k)
natToBin (S (j + j)) | Even = BI (natToBin j)
natToBin (S (S (j + j))) | Odd ?= BO (natToBin (S j))
```

The problem with the Odd case is the same as in the definition of
`parity`

, and the proof proceeds in the same way:

```
natToBin_lemma_1 = proof {
intro;
intro;
rewrite sym (plusSuccRightSucc j j);
trivial;
}
```

To finish, we’ll implement a main program which reads an integer from the user and outputs it in binary.

```
main : IO ()
main = do putStr "Enter a number: "
x <- getLine
print (natToBin (fromInteger (cast x)))
```

For this to work, of course, we need a `Show`

implementation for
`Binary n`

:

```
Show (Binary n) where
show (BO x) = show x ++ "0"
show (BI x) = show x ++ "1"
show BEnd = ""
```

### Interactive Editing¶

By now, we have seen several examples of how Idris’ dependent type
system can give extra confidence in a function’s correctness by giving
a more precise description of its intended behaviour in its *type*. We
have also seen an example of how the type system can help with EDSL
development by allowing a programmer to describe the type system of an
object language. However, precise types give us more than verification
of programs — we can also exploit types to help write programs which
are *correct by construction*.

The Idris REPL provides several commands for inspecting and modifying parts of programs, based on their types, such as case splitting on a pattern variable, inspecting the type of a hole, and even a basic proof search mechanism. In this section, we explain how these features can be exploited by a text editor, and specifically how to do so in Vim. An interactive mode for Emacs is also available.

#### Editing at the REPL¶

The REPL provides a number of commands, which we will describe shortly, which generate new program fragments based on the currently loaded module. These take the general form:

```
:command [line number] [name]
```

That is, each command acts on a specific source line, at a specific
name, and outputs a new program fragment. Each command has an
alternative form, which *updates* the source file in-place:

```
:command! [line number] [name]
```

When the REPL is loaded, it also starts a background process which
accepts and responds to REPL commands, using `idris --client`

. For
example, if we have a REPL running elsewhere, we can execute commands
such as:

```
$ idris --client ':t plus'
Prelude.Nat.plus : Nat -> Nat -> Nat
$ idris --client '2+2'
4 : Integer
```

A text editor can take advantage of this, along with the editing commands, in order to provide interactive editing support.

#### Editing Commands¶

##### :addclause¶

The `:addclause n f`

command, abbreviated `:ac n f`

, creates a
template definition for the function named `f`

declared on line
`n`

. For example, if the code beginning on line 94 contains:

```
vzipWith : (a -> b -> c) ->
Vect n a -> Vect n b -> Vect n c
```

then `:ac 94 vzipWith`

will give:

```
vzipWith f xs ys = ?vzipWith_rhs
```

The names are chosen according to hints which may be given by a programmer, and then made unique by the machine by adding a digit if necessary. Hints can be given as follows:

```
%name Vect xs, ys, zs, ws
```

This declares that any names generated for types in the `Vect`

family
should be chosen in the order `xs`

, `ys`

, `zs`

, `ws`

.

##### :casesplit¶

The `:casesplit n x`

command, abbreviated `:cs n x`

, splits the
pattern variable `x`

on line `n`

into the various pattern forms it
may take, removing any cases which are impossible due to unification
errors. For example, if the code beginning on line 94 is:

```
vzipWith : (a -> b -> c) ->
Vect n a -> Vect n b -> Vect n c
vzipWith f xs ys = ?vzipWith_rhs
```

then `:cs 96 xs`

will give:

```
vzipWith f [] ys = ?vzipWith_rhs_1
vzipWith f (x :: xs) ys = ?vzipWith_rhs_2
```

That is, the pattern variable `xs`

has been split into the two
possible cases `[]`

and `x :: xs`

. Again, the names are chosen
according to the same heuristic. If we update the file (using
`:cs!`

) then case split on `ys`

on the same line, we get:

```
vzipWith f [] [] = ?vzipWith_rhs_3
```

That is, the pattern variable `ys`

has been split into one case
`[]`

, Idris having noticed that the other possible case ```
y ::
ys
```

would lead to a unification error.

##### :addmissing¶

The `:addmissing n f`

command, abbreviated `:am n f`

, adds the
clauses which are required to make the function `f`

on line `n`

cover all inputs. For example, if the code beginning on line 94 is:

```
vzipWith : (a -> b -> c) ->
Vect n a -> Vect n b -> Vect n c
vzipWith f [] [] = ?vzipWith_rhs_1
```

then `:am 96 vzipWith`

gives:

```
vzipWith f (x :: xs) (y :: ys) = ?vzipWith_rhs_2
```

That is, it notices that there are no cases for empty vectors, generates the required clauses, and eliminates the clauses which would lead to unification errors.

##### :proofsearch¶

The `:proofsearch n f`

command, abbreviated `:ps n f`

, attempts to
find a value for the hole `f`

on line `n`

by proof search,
trying values of local variables, recursive calls and constructors of
the required family. Optionally, it can take a list of *hints*, which
are functions it can try applying to solve the hole. For
example, if the code beginning on line 94 is:

```
vzipWith : (a -> b -> c) ->
Vect n a -> Vect n b -> Vect n c
vzipWith f [] [] = ?vzipWith_rhs_1
vzipWith f (x :: xs) (y :: ys) = ?vzipWith_rhs_2
```

then `:ps 96 vzipWith_rhs_1`

will give

```
[]
```

This works because it is searching for a `Vect`

of length 0, of
which the empty vector is the only possibility. Similarly, and perhaps
surprisingly, there is only one possibility if we try to solve ```
:ps
97 vzipWith_rhs_2
```

:

```
f x y :: (vzipWith f xs ys)
```

This works because `vzipWith`

has a precise enough type: The
resulting vector has to be non-empty (a `::`

); the first element
must have type `c`

and the only way to get this is to apply `f`

to
`x`

and `y`

; finally, the tail of the vector can only be built
recursively.

##### :makewith¶

The `:makewith n f`

command, abbreviated `:mw n f`

, adds a
`with`

to a pattern clause. For example, recall `parity`

. If line
10 is:

```
parity (S k) = ?parity_rhs
```

then `:mw 10 parity`

will give:

```
parity (S k) with (_)
parity (S k) | with_pat = ?parity_rhs
```

If we then fill in the placeholder `_`

with `parity k`

and case
split on `with_pat`

using `:cs 11 with_pat`

we get the following
patterns:

```
parity (S (plus n n)) | even = ?parity_rhs_1
parity (S (S (plus n n))) | odd = ?parity_rhs_2
```

Note that case splitting has normalised the patterns here (giving
`plus`

rather than `+`

). In any case, we see that using
interactive editing significantly simplifies the implementation of
dependent pattern matching by showing a programmer exactly what the
valid patterns are.

#### Interactive Editing in Vim¶

The editor mode for Vim provides syntax highlighting, indentation and interactive editing support using the commands described above. Interactive editing is achieved using the following editor commands, each of which update the buffer directly:

`\d`

adds a template definition for the name declared on the- current line (using
`:addclause`

).

`\c`

case splits the variable at the cursor (using`:casesplit`

).

`\m`

adds the missing cases for the name at the cursor (using`:addmissing`

).

`\w`

adds a`with`

clause (using`:makewith`

).`\o`

invokes a proof search to solve the hole under the- cursor (using
`:proofsearch`

).

`\p`

invokes a proof search with additional hints to solve the- hole under the cursor (using
`:proofsearch`

).

There are also commands to invoke the type checker and evaluator:

`\t`

displays the type of the (globally visible) name under the- cursor. In the case of a hole, this displays the context and the expected type.

`\e`

prompts for an expression to evaluate.`\r`

reloads and type checks the buffer.

Corresponding commands are also available in the Emacs mode. Support
for other editors can be added in a relatively straightforward manner
by using `idris –client`

.

### Syntax Extensions¶

Idris supports the implementation of *Embedded Domain Specific
Languages* (EDSLs) in several ways [1]. One way, as we have already
seen, is through extending `do`

notation. Another important way is
to allow extension of the core syntax. In this section we describe two
ways of extending the syntax: `syntax`

rules and `dsl`

notation.

`syntax`

rules¶

We have seen `if...then...else`

expressions, but these are not built
in. Instead, we can define a function in the prelude as follows (we
have already seen this function in Section Laziness):

```
ifThenElse : (x:Bool) -> Lazy a -> Lazy a -> a;
ifThenElse True t e = t;
ifThenElse False t e = e;
```

and then extend the core syntax with a `syntax`

declaration:

```
syntax if [test] then [t] else [e] = ifThenElse test t e;
```

The left hand side of a `syntax`

declaration describes the syntax
rule, and the right hand side describes its expansion. The syntax rule
itself consists of:

**Keywords**— here,`if`

,`then`

and`else`

, which must be valid identifiers.**Non-terminals**— included in square brackets,`[test]`

,`[t]`

and`[e]`

here, which stand for arbitrary expressions. To avoid parsing ambiguities, these expressions cannot use syntax extensions at the top level (though they can be used in parentheses).**Names**— included in braces, which stand for names which may be bound on the right hand side.**Symbols**— included in quotations marks, e.g.`":="`

. This can also be used to include reserved words in syntax rules, such as`"let"`

or`"in"`

.

The limitations on the form of a syntax rule are that it must include at least one symbol or keyword, and there must be no repeated variables standing for non-terminals. Any expression can be used, but if there are two non-terminals in a row in a rule, only simple expressions may be used (that is, variables, constants, or bracketed expressions). Rules can use previously defined rules, but may not be recursive. The following syntax extensions would therefore be valid:

```
syntax [var] ":=" [val] = Assign var val;
syntax [test] "?" [t] ":" [e] = if test then t else e;
syntax select [x] from [t] "where" [w] = SelectWhere x t w;
syntax select [x] from [t] = Select x t;
```

Syntax macros can be further restricted to apply only in patterns (i.e.
only on the left hand side of a pattern match clause) or only in terms
(i.e. everywhere but the left hand side of a pattern match clause) by
being marked as `pattern`

or `term`

syntax rules. For example, we
might define an interval as follows, with a static check that the lower
bound is below the upper bound using `so`

:

```
data Interval : Type where
MkInterval : (lower : Double) -> (upper : Double) ->
So (lower < upper) -> Interval
```

We can define a syntax which, in patterns, always matches `Oh`

for
the proof argument, and in terms requires a proof term to be provided:

```
pattern syntax "[" [x] "..." [y] "]" = MkInterval x y Oh
term syntax "[" [x] "..." [y] "]" = MkInterval x y ?bounds_lemma
```

In terms, the syntax `[x...y]`

will generate a proof obligation
`bounds_lemma`

(possibly renamed).

Finally, syntax rules may be used to introduce alternative binding
forms. For example, a `for`

loop binds a variable on each iteration:

```
syntax for {x} "in" [xs] ":" [body] = forLoop xs (\x => body)
main : IO ()
main = do for x in [1..10]:
putStrLn ("Number " ++ show x)
putStrLn "Done!"
```

Note that we have used the `{x}`

form to state that `x`

represents
a bound variable, substituted on the right hand side. We have also put
`in`

in quotation marks since it is already a reserved word.

`dsl`

notation¶

The well-typed interpreter in Section Example: The Well-Typed Interpreter is a simple
example of a common programming pattern with dependent types. Namely:
describe an *object language* and its type system with dependent types
to guarantee that only well-typed programs can be represented, then
program using that representation. Using this approach we can, for
example, write programs for serialising binary data [2] or running
concurrent processes safely [3].

Unfortunately, the form of object language programs makes it rather
hard to program this way in practice. Recall the factorial program in
`Expr`

for example:

```
fact : Expr G (TyFun TyInt TyInt)
fact = Lam (If (Op (==) (Var Stop) (Val 0))
(Val 1) (Op (*) (App fact (Op (-) (Var Stop) (Val 1)))
(Var Stop)))
```

Since this is a particularly useful pattern, Idris provides syntax overloading [1] to make it easier to program in such object languages:

```
mkLam : TTName -> Expr (t::g) t' -> Expr g (TyFun t t')
mkLam _ body = Lam body
dsl expr
variable = Var
index_first = Stop
index_next = Pop
lambda = mkLam
```

A `dsl`

block describes how each syntactic construct is represented
in an object language. Here, in the `expr`

language, any variable is
translated to the `Var`

constructor, using `Pop`

and `Stop`

to
construct the de Bruijn index (i.e., to count how many bindings since
the variable itself was bound); and any lambda is translated to a
`Lam`

constructor. The `mkLam`

function simply ignores its first
argument, which is the name that the user chose for the variable. It
is also possible to overload `let`

and dependent function syntax
(`pi`

) in this way. We can now write `fact`

as follows:

```
fact : Expr G (TyFun TyInt TyInt)
fact = expr (\x => If (Op (==) x (Val 0))
(Val 1) (Op (*) (app fact (Op (-) x (Val 1))) x))
```

In this new version, `expr`

declares that the next expression will
be overloaded. We can take this further, using idiom brackets, by
declaring:

```
(<*>) : (f : Lazy (Expr G (TyFun a t))) -> Expr G a -> Expr G t
(<*>) f a = App f a
pure : Expr G a -> Expr G a
pure = id
```

Note that there is no need for these to be part of an implementation of
`Applicative`

, since idiom bracket notation translates directly to
the names `<*>`

and `pure`

, and ad-hoc type-directed overloading
is allowed. We can now say:

```
fact : Expr G (TyFun TyInt TyInt)
fact = expr (\x => If (Op (==) x (Val 0))
(Val 1) (Op (*) [| fact (Op (-) x (Val 1)) |] x))
```

With some more ad-hoc overloading and use of interfaces, and a new syntax rule, we can even go as far as:

```
syntax "IF" [x] "THEN" [t] "ELSE" [e] = If x t e
fact : Expr G (TyFun TyInt TyInt)
fact = expr (\x => IF x == 0 THEN 1 ELSE [| fact (x - 1) |] * x)
```

[1] | (1, 2) Edwin Brady and Kevin Hammond. 2012. Resource-Safe systems
programming with embedded domain specific languages. In
Proceedings of the 14th international conference on Practical
Aspects of Declarative Languages (PADL‘12), Claudio Russo and
Neng-Fa Zhou (Eds.). Springer-Verlag, Berlin, Heidelberg,
242-257. DOI=10.1007/978-3-642-27694-1_18
http://dx.doi.org/10.1007/978-3-642-27694-1_18 |

[2] | Edwin C. Brady. 2011. IDRIS —: systems programming meets full dependent types. In Proceedings of the 5th ACM workshop on Programming languages meets program verification (PLPV ‘11). ACM, New York, NY, USA, 43-54. DOI=10.1145/1929529.1929536 http://doi.acm.org/10.1145/1929529.1929536 |

[3] | Edwin Brady and Kevin Hammond. 2010. Correct-by-Construction Concurrency: Using Dependent Types to Verify Implementations of Effectful Resource Usage Protocols. Fundam. Inf. 102, 2 (April 2010), 145-176. http://dl.acm.org/citation.cfm?id=1883636 |

### Miscellany¶

In this section we discuss a variety of additional features:

- auto, implicit, and default arguments;
- literate programming;
- interfacing with external libraries through the foreign function;
- interface;
- type providers;
- code generation; and
- the universe hierarchy.

#### Implicit arguments¶

We have already seen implicit arguments, which allows arguments to be omitted when they can be inferred by the type checker, e.g.

```
index : {a:Type} -> {n:Nat} -> Fin n -> Vect n a -> a
```

##### Auto implicit arguments¶

In other situations, it may be possible to infer arguments not by type
checking but by searching the context for an appropriate value, or
constructing a proof. For example, the following definition of `head`

which requires a proof that the list is non-empty:

```
isCons : List a -> Bool
isCons [] = False
isCons (x :: xs) = True
head : (xs : List a) -> (isCons xs = True) -> a
head (x :: xs) _ = x
```

If the list is statically known to be non-empty, either because its
value is known or because a proof already exists in the context, the
proof can be constructed automatically. Auto implicit arguments allow
this to happen silently. We define `head`

as follows:

```
head : (xs : List a) -> {auto p : isCons xs = True} -> a
head (x :: xs) = x
```

The `auto`

annotation on the implicit argument means that Idris
will attempt to fill in the implicit argument by searching for a value
of the appropriate type. It will try the following, in order:

- Local variables, i.e. names bound in pattern matches or
`let`

bindings, with exactly the right type. - The constructors of the required type. If they have arguments, it will search recursively up to a maximum depth of 100.
- Local variables with function types, searching recursively for the arguments.
- Any function with the appropriate return type which is marked with the
`%hint`

annotation.

In the case that a proof is not found, it can be provided explicitly as normal:

```
head xs {p = ?headProof}
```

##### Default implicit arguments¶

Besides having Idris automatically find a value of a given type, sometimes we
want to have an implicit argument with a specific default value. In Idris, we can
do this using the `default`

annotation. While this is primarily intended to assist
in automatically constructing a proof where auto fails, or finds an unhelpful value,
it might be easier to first consider a simpler case, not involving proofs.

If we want to compute the n’th fibonacci number (and defining the 0th fibonacci number as 0), we could write:

```
fibonacci : {default 0 lag : Nat} -> {default 1 lead : Nat} -> (n : Nat) -> Nat
fibonacci {lag} Z = lag
fibonacci {lag} {lead} (S n) = fibonacci {lag=lead} {lead=lag+lead} n
```

After this definition, `fibonacci 5`

is equivalent to `fibonacci {lag=0} {lead=1} 5`

,
and will return the 5th fibonacci number. Note that while this works, this is not the
intended use of the `default`

annotation. It is included here for illustrative purposes
only. Usually, `default`

is used to provide things like a custom proof search script.

#### Implicit conversions¶

Idris supports the creation of *implicit conversions*, which allow
automatic conversion of values from one type to another when required to
make a term type correct. This is intended to increase convenience and
reduce verbosity. A contrived but simple example is the following:

```
implicit intString : Int -> String
intString = show
test : Int -> String
test x = "Number " ++ x
```

In general, we cannot append an `Int`

to a `String`

, but the
implicit conversion function `intString`

can convert `x`

to a
`String`

, so the definition of `test`

is type correct. An implicit
conversion is implemented just like any other function, but given the
`implicit`

modifier, and restricted to one explicit argument.

Only one implicit conversion will be applied at a time. That is, implicit conversions cannot be chained. Implicit conversions of simple types, as above, are however discouraged! More commonly, an implicit conversion would be used to reduce verbosity in an embedded domain specific language, or to hide details of a proof. Such examples are beyond the scope of this tutorial.

#### Literate programming¶

Like Haskell, Idris supports *literate* programming. If a file has
an extension of `.lidr`

then it is assumed to be a literate file. In
literate programs, everything is assumed to be a comment unless the line
begins with a greater than sign `>`

, for example:

```
> module literate
This is a comment. The main program is below
> main : IO ()
> main = putStrLn "Hello literate world!\n"
```

An additional restriction is that there must be a blank line between a
program line (beginning with `>`

) and a comment line (beginning with
any other character).

#### Foreign function calls¶

For practical programming, it is often necessary to be able to use
external libraries, particularly for interfacing with the operating
system, file system, networking, *et cetera*. Idris provides a
lightweight foreign function interface for achieving this, as part of
the prelude. For this, we assume a certain amount of knowledge of C and
the `gcc`

compiler. First, we define a datatype which describes the
external types we can handle:

```
data FTy = FInt | FFloat | FChar | FString | FPtr | FUnit
```

Each of these corresponds directly to a C type. Respectively: `int`

,
`double`

, `char`

, `char*`

, `void*`

and `void`

. There is also a
translation to a concrete Idris type, described by the following
function:

```
interpFTy : FTy -> Type
interpFTy FInt = Int
interpFTy FFloat = Double
interpFTy FChar = Char
interpFTy FString = String
interpFTy FPtr = Ptr
interpFTy FUnit = ()
```

A foreign function is described by a list of input types and a return type, which can then be converted to an Idris type:

```
ForeignTy : (xs:List FTy) -> (t:FTy) -> Type
```

A foreign function is assumed to be impure, so `ForeignTy`

builds an
`IO`

type, for example:

```
Idris> ForeignTy [FInt, FString] FString
Int -> String -> IO String : Type
Idris> ForeignTy [FInt, FString] FUnit
Int -> String -> IO () : Type
```

We build a call to a foreign function by giving the name of the
function, a list of argument types and the return type. The built in
construct `mkForeign`

converts this description to a function callable
by Idris:

```
data Foreign : Type -> Type where
FFun : String -> (xs:List FTy) -> (t:FTy) ->
Foreign (ForeignTy xs t)
mkForeign : Foreign x -> x
```

Note that the compiler expects `mkForeign`

to be fully applied to
build a complete foreign function call. For example, the `putStr`

function is implemented as follows, as a call to an external function
`putStr`

defined in the run-time system:

```
putStr : String -> IO ()
putStr x = mkForeign (FFun "putStr" [FString] FUnit) x
```

##### Include and linker directives¶

Foreign function calls are translated directly to calls to C functions, with appropriate conversion between the Idris representation of a value and the C representation. Often this will require extra libraries to be linked in, or extra header and object files. This is made possible through the following directives:

`%lib target x`

— include the`libx`

library. If the target is`C`

this is equivalent to passing the`-lx`

option to`gcc`

. If the target is Java the library will be interpreted as a`groupId:artifactId:packaging:version`

dependency coordinate for maven.`%include target x`

— use the header file or import`x`

for the given back end target.`%link target x.o`

— link with the object file`x.o`

when using the given back end target.`%dynamic x.so`

— dynamically link the interpreter with the shared object`x.so`

.

##### Testing foreign function calls¶

Normally, the Idris interpreter (used for typechecking and at the REPL)
will not perform IO actions. Additionally, as it neither generates C
code nor compiles to machine code, the `%lib`

, `%include`

and
`%link`

directives have no effect. IO actions and FFI calls can be
tested using the special REPL command `:x EXPR`

, and C libraries can
be dynamically loaded in the interpreter by using the `:dynamic`

command or the `%dynamic`

directive. For example:

```
Idris> :dynamic libm.so
Idris> :x unsafePerformIO ((mkForeign (FFun "sin" [FFloat] FFloat)) 1.6)
0.9995736030415051 : Double
```

#### Type Providers¶

Idris type providers, inspired by F#’s type providers, are a means of making our types be “about” something in the world outside of Idris. For example, given a type that represents a database schema and a query that is checked against it, a type provider could read the schema of a real database during type checking.

Idris type providers use the ordinary execution semantics of Idris to run an IO action and extract the result. This result is then saved as a constant in the compiled code. It can be a type, in which case it is used like any other type, or it can be a value, in which case it can be used as any other value, including as an index in types.

Type providers are still an experimental extension. To enable the
extension, use the `%language`

directive:

```
%language TypeProviders
```

A provider `p`

for some type `t`

is simply an expression of type
`IO (Provider t)`

. The `%provide`

directive causes the type checker
to execute the action and bind the result to a name. This is perhaps
best illustrated with a simple example. The type provider `fromFile`

reads a text file. If the file consists of the string `Int`

, then the
type `Int`

will be provided. Otherwise, it will provide the type
`Nat`

.

```
strToType : String -> Type
strToType "Int" = Int
strToType _ = Nat
fromFile : String -> IO (Provider Type)
fromFile fname = do Right str <- readFile fname
| Left err => pure (Provide Void)
pure (Provide (strToType (trim str)))
```

We then use the `%provide`

directive:

```
%provide (T1 : Type) with fromFile "theType"
foo : T1
foo = 2
```

If the file named `theType`

consists of the word `Int`

, then `foo`

will be an `Int`

. Otherwise, it will be a `Nat`

. When Idris
encounters the directive, it first checks that the provider expression
`fromFile theType`

has type `IO (Provider Type)`

. Next, it executes
the provider. If the result is `Provide t`

, then `T1`

is defined as
`t`

. Otherwise, the result is an error.

Our datatype `Provider t`

has the following definition:

```
data Provider a = Error String
| Provide a
```

We have already seen the `Provide`

constructor. The `Error`

constructor allows type providers to return useful error messages. The
example in this section was purposefully simple. More complex type
provider implementations, including a statically-checked SQLite binding,
are available in an external collection [1].

#### C Target¶

The default target of Idris is C. Compiling via:

```
$ idris hello.idr -o hello
```

is equivalent to:

```
$ idris --codegen C hello.idr -o hello
```

When the command above is used, a temporary C source is generated, which
is then compiled into an executable named `hello`

.

In order to view the generated C code, compile via:

```
$ idris hello.idr -S -o hello.c
```

To turn optimisations on, use the `%flag C`

pragma within the code, as
is shown below:

```
module Main
%flag C "-O3"
factorial : Int -> Int
factorial 0 = 1
factorial n = n * (factorial (n-1))
main : IO ()
main = do
putStrLn $ show $ factorial 3
```

To compile the generated C with debugging information e.g. to use
`gdb`

to debug segmentation faults in Idris programs, use the
`%flag C`

pragma to include debugging symbols, as is shown below:

```
%flag C "-g"
```

#### JavaScript Target¶

Idris is capable of producing *JavaScript* code that can be run in a
browser as well as in the *NodeJS* environment or alike. One can use the
FFI to communicate with the *JavaScript* ecosystem.

##### Code Generation¶

Code generation is split into two separate targets. To generate code that is tailored for running in the browser issue the following command:

```
$ idris --codegen javascript hello.idr -o hello.js
```

The resulting file can be embedded into your HTML just like any other
*JavaScript* code.

Generating code for *NodeJS* is slightly different. Idris outputs a
*JavaScript* file that can be directly executed via `node`

.

```
$ idris --codegen node hello.idr -o hello
$ ./hello
Hello world
```

Take into consideration that the *JavaScript* code generator is using
`console.log`

to write text to `stdout`

, this means that it will
automatically add a newline to the end of each string. This behaviour
does not show up in the *NodeJS* code generator.

##### Using the FFI¶

To write a useful application we need to communicate with the outside
world. Maybe we want to manipulate the DOM or send an Ajax request. For
this task we can use the FFI. Since most *JavaScript* APIs demand
callbacks we need to extend the FFI so we can pass functions as
arguments.

The *JavaScript* FFI works a little bit differently than the regular
FFI. It uses positional arguments to directly insert our arguments into
a piece of *JavaScript* code.

One could use the primitive addition of *JavaScript* like so:

```
module Main
primPlus : Int -> Int -> IO Int
primPlus a b = mkForeign (FFun "%0 + %1" [FInt, FInt] FInt) a b
main : IO ()
main = do
a <- primPlus 1 1
b <- primPlus 1 2
print (a, b)
```

Notice that the `%n`

notation qualifies the position of the `n`

-th
argument given to our foreign function starting from 0. When you need a
percent sign rather than a position simply use `%%`

instead.

Passing functions to a foreign function is very similar. Let’s assume
that we want to call the following function from the *JavaScript* world:

```
function twice(f, x) {
return f(f(x));
}
```

We obviously need to pass a function `f`

here (we can infer it from
the way we use `f`

in `twice`

, it would be more obvious if
*JavaScript* had types).

The *JavaScript* FFI is able to understand functions as arguments when
you give it something of type `FFunction`

. The following example code
calls `twice`

in *JavaScript* and returns the result to our Idris
program:

```
module Main
twice : (Int -> Int) -> Int -> IO Int
twice f x = mkForeign (
FFun "twice(%0,%1)" [FFunction FInt FInt, FInt] FInt
) f x
main : IO ()
main = do
a <- twice (+1) 1
print a
```

The program outputs `3`

, just like we expected.

##### Including external *JavaScript* files¶

Whenever one is working with *JavaScript* one might want to include
external libraries or just some functions that she or he wants to call
via FFI which are stored in external files. The *JavaScript* and
*NodeJS* code generators understand the `%include`

directive. Keep in
mind that *JavaScript* and *NodeJS* are handled as different code
generators, therefore you will have to state which one you want to
target. This means that you can include different files for *JavaScript*
and *NodeJS* in the same Idris source file.

So whenever you want to add an external *JavaScript* file you can do
this like so:

For *NodeJS*:

```
%include Node "path/to/external.js"
```

And for use in the browser:

```
%include JavaScript "path/to/external.js"
```

The given files will be added to the top of the generated code. For library packages you can also use the ipkg objs option to include the js file in the installation, and use:

```
%include Node "package/external.js"
```

The *JavaScript* and *NodeJS* backends of Idris will also lookup for the file
on that location.

##### Including *NodeJS* modules¶

The *NodeJS* code generator can also include modules with the `%lib`

directive.

```
%lib Node "fs"
```

This directive compiles into the following *JavaScript*

```
var fs = require("fs");
```

##### Shrinking down generated *JavaScript*¶

Idris can produce very big chunks of *JavaScript* code. However, the
generated code can be minified using the `closure-compiler`

from
Google. Any other minifier is also suitable but `closure-compiler`

offers advanced compilation that does some aggressive inlining and code
elimination. Idris can take full advantage of this compilation mode
and it’s highly recommended to use it when shipping a *JavaScript*
application written in Idris.

#### Cumulativity¶

Since values can appear in types and *vice versa*, it is natural that
types themselves have types. For example:

```
*universe> :t Nat
Nat : Type
*universe> :t Vect
Vect : Nat -> Type -> Type
```

But what about the type of `Type`

? If we ask Idris it reports:

```
*universe> :t Type
Type : Type 1
```

If `Type`

were its own type, it would lead to an inconsistency due to
Girard’s paradox,
so internally there is a *hierarchy* of types (or *universes*):

```
Type : Type 1 : Type 2 : Type 3 : ...
```

Universes are *cumulative*, that is, if `x : Type n`

we can also have
that `x : Type m`

, as long as `n < m`

. The typechecker generates
such universe constraints and reports an error if any inconsistencies
are found. Ordinarily, a programmer does not need to worry about this,
but it does prevent (contrived) programs such as the following:

```
myid : (a : Type) -> a -> a
myid _ x = x
idid : (a : Type) -> a -> a
idid = myid _ myid
```

The application of `myid`

to itself leads to a cycle in the universe
hierarchy — `myid`

’s first argument is a `Type`

, which cannot be
at a lower level than required if it is applied to itself.

[1] | https://github.com/david-christiansen/idris-type-providers |

### Further Reading¶

Further information about Idris programming, and programming with dependent types in general, can be obtained from various sources:

- The Idris web site (http://www.idris-lang.org/) and by asking questions on the mailing list.
- The IRC channel
`#idris`

, on chat.freenode.net. - The wiki (https://github.com/idris-lang/Idris-dev/wiki/) has further
- user provided information, in particular:

- Examining the prelude and exploring the
`samples`

in the - distribution. The Idris source can be found online at: https://github.com/idris-lang/Idris-dev.

- Examining the prelude and exploring the
- Existing projects on the
`Idris Hackers`

web space: http://idris-hackers.github.io.

[1] | Edwin Brady and Kevin Hammond. 2012. Resource-Safe systems programming with embedded domain specific languages. In Proceedings of the 14th international conference on Practical Aspects of Declarative Languages (PADL‘12), Claudio Russo and Neng-Fa Zhou (Eds.). Springer-Verlag, Berlin, Heidelberg, 242-257. DOI=10.1007/978-3-642-27694-1_18 http://dx.doi.org/10.1007/978-3-642-27694-1_18 |

[2] | Edwin C. Brady. 2011. IDRIS —: systems programming meets full dependent types. In Proceedings of the 5th ACM workshop on Programming languages meets program verification (PLPV ‘11). ACM, New York, NY, USA, 43-54. DOI=10.1145/1929529.1929536 http://doi.acm.org/10.1145/1929529.1929536 |

[3] | Edwin C. Brady and Kevin Hammond. 2010. Scrapping your inefficient engine: using partial evaluation to improve domain-specific language implementation. In Proceedings of the 15th ACM SIGPLAN international conference on Functional programming (ICFP ‘10). ACM, New York, NY, USA, 297-308. DOI=10.1145/1863543.1863587 http://doi.acm.org/10.1145/1863543.1863587 |

## Frequently Asked Questions¶

### What are the differences between Agda and Idris?¶

Like Idris, Agda is a functional language with dependent types, supporting dependent pattern matching. Both can be used for writing programs and proofs. However, Idris has been designed from the start to emphasise general purpose programming rather than theorem proving. As such, it supports interoperability with systems libraries and C programs, and language constructs for domain specific language implementation. It also includes higher level programming constructs such as interfaces (similar to type classes) and do notation.

Idris supports multiple back ends (C and JavaScript by default, with the ability to add more via plugins) and has a reference run time system, written in C, with a garbage collector and built-in message passing concurrency.

### Is Idris production ready?¶

Idris is primarily a research tool for exploring the possibilities of software development with dependent types, meaning that the primary goal is not (yet) to make a system which could be used in production. As such, there are a few rough corners, and lots of missing libraries. Nobody is working on Idris full time, and we don’t have the resources at the moment to polish the system on our own. Therefore, we don’t recommend building your business around it!

Having said that, contributions which help towards making Idris suitable for use in production would be very welcome - this includes (but is not limited to) extra library support, polishing the run-time system (and ensuring it is robust), providing and maintaining a JVM back end, etc.

### Why does Idris use eager evaluation rather than lazy?¶

Idris uses eager evaluation for more predictable performance, in particular
because one of the longer term goals is to be able to write efficient and
verified low level code such as device drivers and network infrastructure.
Furthermore, the Idris type system allows us to state precisely the type
of each value, and therefore the run-time form of each value. In a lazy
language, consider a value of type `Int`

:

```
thing : Int
```

What is the representation of `thing`

at run-time? Is it a bit pattern
representing an integer, or is it a pointer to some code which will compute
an integer? In Idris, we have decided that we would like to make this
distinction precise, in the type:

```
thing_val : Int
thing_comp : Lazy Int
```

Here, it is clear from the type that `thing_val`

is guaranteed to be a
concrete `Int`

, whereas `thing_comp`

is a computation which will produce an
`Int`

.

### How can I make lazy control structures?¶

You can make control structures using the special Lazy type. For
example, `if...then...else...`

in Idris expands to an application of
a function named `ifThenElse`

. The default implementation for
Booleans is defined as follows in the library:

```
ifThenElse : Bool -> (t : Lazy a) -> (e : Lazy a) -> a
ifThenElse True t e = t
ifThenElse False t e = e
```

The type `Lazy a`

for `t`

and `e`

indicates that those arguments will
only be evaluated if they are used, that is, they are evaluated lazily.

### Evaluation at the REPL doesn’t behave as I expect. What’s going on?¶

Being a fully dependently typed language, Idris has two phases where it evaluates things, compile-time and run-time. At compile-time it will only evaluate things which it knows to be total (i.e. terminating and covering all possible inputs) in order to keep type checking decidable. The compile-time evaluator is part of the Idris kernel, and is implemented in Haskell using a HOAS (higher order abstract syntax) style representation of values. Since everything is known to have a normal form here, the evaluation strategy doesn’t actually matter because either way it will get the same answer, and in practice it will do whatever the Haskell run-time system chooses to do.

The REPL, for convenience, uses the compile-time notion of evaluation. As well as being easier to implement (because we have the evaluator available) this can be very useful to show how terms evaluate in the type checker. So you can see the difference between:

```
Idris> \n, m => (S n) + m
\n => \m => S (plus n m) : Nat -> Nat -> Nat
Idris> \n, m => n + (S m)
\n => \m => plus n (S m) : Nat -> Nat -> Nat
```

### Why can’t I use a function with no arguments in a type?¶

If you use a name in a type which begins with a lower case letter, and which is not applied to any arguments, then Idris will treat it as an implicitly bound argument. For example:

```
append : Vect n ty -> Vect m ty -> Vect (n + m) ty
```

Here, `n`

, `m`

, and `ty`

are implicitly bound. This rule applies even
if there are functions defined elsewhere with any of these names. For example,
you may also have:

```
ty : Type
ty = String
```

Even in this case, `ty`

is still considered implicitly bound in the definition
of `append`

, rather than making the type of `append`

equivalent to…

```
append : Vect n String -> Vect m String -> Vect (n + m) String
```

…which is probably not what was intended! The reason for this rule is so
that it is clear just from looking at the type of `append`

, and no other
context, what the implicitly bound names are.

If you want to use an unapplied name in a type, you have two options. You
can either explicitly qualify it, for example, if `ty`

is defined in the
namespace `Main`

you can do the following:

```
append : Vect n Main.ty -> Vect m Main.ty -> Vect (n + m) Main.ty
```

Alternatively, you can use a name which does not begin with a lower case letter, which will never be implicitly bound:

```
Ty : Type
Ty = String
append : Vect n Ty -> Vect m Ty -> Vect (n + m) Ty
```

As a convention, if a name is intended to be used as a type synonym, it is best for it to begin with a capital letter to avoid this restriction.

### I have an obviously terminating program, but Idris says it possibly isn’t total. Why is that?¶

Idris can’t decide in general whether a program is terminating due to the undecidability of the Halting Problem. It is possible, however, to identify some programs which are definitely terminating. Idris does this using “size change termination” which looks for recursive paths from a function back to itself. On such a path, there must be at least one argument which converges to a base case.

- Mutually recursive functions are supported
- However, all functions on the path must be fully applied. In particular, higher order applications are not supported
- Idris identifies arguments which converge to a base case by looking for
recursive calls to syntactically smaller arguments of inputs. e.g.
`k`

is syntactically smaller than`S (S k)`

because`k`

is a subterm of`S (S k)`

, but`(k, k)`

is not syntactically smaller than`(S k, S k)`

.

If you have a function which you believe to be terminating, but Idris does
not, you can either restructure the program, or use the `assert_total`

function.

### When will Idris be self-hosting?¶

It’s not a priority, though not a bad idea in the long run. It would be a worthwhile effort in the short term to implement libraries to support self-hosting, such as a good parsing library.

### Does Idris have universe polymorphism? What is the type of `Type`

?¶

Rather than universe polymorphism, Idris has a cumulative hierarchy of
universes; `Type : Type 1`

, `Type 1 : Type 2`

, etc.
Cumulativity means that if `x : Type n`

and `n <= m`

, then
`x : Type m`

. Universe levels are always inferred by Idris, and
cannot be specified explicitly. The REPL command `:type Type 1`

will
result in an error, as will attempting to specify the universe level
of any type.

### Why does Idris use `Double`

instead of `Float64`

?¶

Historically the C language and many other languages have used the
names `Float`

and `Double`

to represent floating point numbers of
size 32 and 64 respectively. Newer languages such as Rust and Julia
have begun to follow the naming scheme described in IEEE Standard for
Floating-Point Arithmetic (IEEE 754). This describes
single and double precision numbers as `Float32`

and `Float64`

;
the size is described in the type name.

Due to developer familiarity with the older naming convention, and
choice by the developers of Idris, Idris uses the C style convention.
That is, the name `Double`

is used to describe double precision
numbers, and Idris does not support 32 bit floats at present.

### What is -ffreestanding?¶

The freestanding flag is used to build Idris binaries which have their libs and compiler in a relative path. This is useful for building binaries where the install directory is unknown at build time. When passing this flag, the IDRIS_LIB_DIR environment variable needs to be set to the path where the Idris libs reside relative to the idris executable. The IDRIS_TOOLCHAIN_DIR environment variable is optional, if that is set, Idris will use that path to find the C compiler. For example:

```
IDRIS_LIB_DIR="./libs" \
IDRIS_TOOLCHAIN_DIR="./mingw/bin" \
CABALFLAGS="-fffi -ffreestanding -frelease" \
make
```

### What does the name “Idris” mean?¶

British people of a certain age may be familiar with this singing dragon. If that doesn’t help, maybe you can invent a suitable acronym :-) .

### Will there be support for Unicode characters for operators?¶

There are several reasons why we should not support Unicode operators:

- It’s hard to type (this is important if you’re using someone else’s code, for example). Various editors have their own input methods, but you have to know what they are.
- Not every piece of software easily supports it. Rendering issues have been noted on some mobile email clients, terminal-based IRC clients, web browsers, etc. There are ways to resolve these rendering issues but they provide a barrier to entry to using Idris.
- Even if we leave it out of the standard library (which we will in any case!) as soon as people start using it in their library code, others have to deal with it.
- Too many characters look too similar. We had enough trouble with confusion between 0 and O without worrying about all the different kinds of colons and brackets.
- There seems to be a tendency to go over the top with use of Unicode. For example, using sharp and flat for delay and force (or is it the other way around?) in Agda seems gratuitous. We don’t want to encourage this sort of thing, when words are often better.

With care, Unicode operators can make things look pretty but so can `lhs2TeX`

.
Perhaps in a few years time things will be different and software will cope
better and it will make sense to revisit this. For now, however, Idris will not
be offering arbitrary Unicode symbols in operators.

This seems like an instance of Wadler’s Law in action.

This answer is based on Edwin Brady’s response in the following pull request.

### Where can I find the community standards for the Idris community?¶

The Idris Community Standards are stated here .

### Where can I find more answers?¶

There is an Unofficial FAQ on the wiki on GitHub which answers more technical questions and may be updated more often.

## Implementing State-aware Systems in Idris: The ST Tutorial¶

A tutorial on implementing state-aware systems using the Control.ST library in Idris.

Note

The documentation for Idris has been published under the Creative
Commons CC0 License. As such to the extent possible under law, *The
Idris Community* has waived all copyright and related or neighbouring
rights to Documentation for Idris.

More information concerning the CC0 can be found online at: http://creativecommons.org/publicdomain/zero/1.0/

### Overview¶

Pure functional languages with dependent types such as Idris support reasoning about programs directly
in the type system, promising that we can *know* a program will run
correctly (i.e. according to the specification in its type) simply
because it compiles.

Realistically, though, software relies on state, and many components rely on state machines. For example, they describe network transport protocols like TCP, and implement event-driven systems and regular expression matching. Furthermore, many fundamental resources like network sockets and files are, implicitly, managed by state machines, in that certain operations are only valid on resources in certain states, and those operations can change the states of the underlying resource. For example, it only makes sense to send a message on a connected network socket, and closing a socket changes its state from “open” to “closed”. State machines can also encode important security properties. For example, in the software which implements an ATM, it’s important that the ATM dispenses cash only when the machine is in a state where a card has been inserted and the PIN verified.

In this tutorial we will introduce the `Control.ST`

library, which is included
with the Idris distribution (currently as part of the `contrib`

package)
and supports programming and reasoning with state and side effects. This
tutorial assumes familiarity with pure programming in Idris, as described in
The Idris Tutorial.
For further background information, the `ST`

library is based on ideas
discussed in Chapter 13 (available as a free sample chapter) and Chapter 14
of Type-Driven Development with Idris.

The `ST`

library allows us to write programs which are composed of multiple
state transition systems. It supports composition in two ways: firstly, we can
use several independently implemented state transition systems at once;
secondly, we can implement one state transition system in terms of others.

#### Introductory example: a data store requiring a login¶

Many software components rely on some form of state, and there may be operations which are only valid in specific states. For example, consider a secure data store in which a user must log in before getting access to some secret data. This system can be in one of two states:

`LoggedIn`

, in which the user is allowed to read the secret`LoggedOut`

, in which the user has no access to the secret

We can provide commands to log in, log out, and read the data, as illustrated in the following diagram:

The `login`

command, if it succeeds, moves the overall system state from
`LoggedOut`

to `LoggedIn`

. The `logout`

command moves the state from
`LoggedIn`

to `LoggedOut`

. Most importantly, the `readSecret`

command
is only valid when the system is in the `LoggedIn`

state.

We routinely use type checkers to ensure that variables and arguments are used
consistently. However, statically checking that operations are performed only
on resources in an appropriate state is not well supported by mainstream type
systems. In the data store example, for example, it’s important to check that
the user is successfully logged in before using `readSecret`

. The
`ST`

library allows us to represent this kind of *protocol* in the type
system, and ensure at *compile-time* that the secret is only read when the
user is logged in.

#### Outline¶

This tutorial starts (Introducing ST: Working with State) by describing how to manipulate
individual states, introduces a data type `STrans`

for describing stateful
functions, and `ST`

which describes top level state transitions.
Next (State Machines in Types) it describes how to represent state machines in
types, and how to define *interfaces* for describing stateful systems.
Then (Composing State Machines) it describes how to compose systems of multiple
state machines. It explains how to implement systems which use several
state machines at once, and how to implement a high level stateful system
in terms of lower level systems.
Finally (Example: Network Socket Programming) we’ll see a specific example of a stateful
API in practice, implementing the POSIX network sockets API.

The `Control.ST`

library is also described in a draft paper by
Edwin Brady, “State Machines All The Way
Down”, available here.
This paper presents many of the examples from this tutorial, and describes
the motivation, design and implementation of the library in more depth.

### Introducing ST: Working with State¶

The `Control.ST`

library provides facilities for creating, reading, writing
and destroying state in Idris functions, and tracking changes of state in
a function’s type. It is based around the concept of *resources*, which are,
essentially, mutable variables, and a dependent type, `STrans`

which tracks
how those resources change when a function runs:

```
STrans : (m : Type -> Type) ->
(resultType : Type) ->
(in_res : Resources) ->
(out_res : resultType -> Resources) ->
Type
```

A value of type `STrans m resultType in_res out_res_fn`

represents a sequence
of actions which can manipulate state. The arguments are:

`m`

, which is an underlying*computation context*in which the actions will be executed. Usually, this will be a generic type with a`Monad`

implementation, but it isn’t necessarily so. In particular, there is no need to understand monads to be able to use`ST`

effectively!`resultType`

, which is the type of the value the sequence will produce`in_res`

, which is a list of*resources*available*before*executing the actions.`out_res`

, which is a list of resources available*after*executing the actions, and may differ depending on the result of the actions.

We can use `STrans`

to describe *state transition systems* in a function’s
type. We’ll come to the definition of `Resources`

shortly, but for the moment
you can consider it an abstract representation of the “state of the world”.
By giving the input resources (`in_res`

) and the output resources
(`out_res`

) we are describing the *preconditions* under which a function
is allowed to execute, and *postconditions* which describe how a function
affects the overall state of the world.

We’ll begin in this section by looking at some small examples of `STrans`

functions, and see how to execute them. We’ll also introduce `ST`

,
a type-level function which allows us to describe the state transitions of
a stateful function concisely.

Type checking the examples

For the examples in this section, and throughout this tutorial,
you’ll need to `import Control.ST`

and add the `contrib`

package by
passing the `-p contrib`

flag to `idris`

.

#### Introductory examples: manipulating `State`

¶

An `STrans`

function explains, in its type, how it affects a collection of
`Resources`

. A resource has a *label* (of type `Var`

), which we use to
refer to the resource throughout the function, and we write the state of a
resource, in the `Resources`

list, in the form `label ::: type`

.

For example, the following function
has a resource `x`

available on input, of type `State Integer`

, and that
resource is still a `State Integer`

on output:

```
increment : (x : Var) -> STrans m () [x ::: State Integer]
(const [x ::: State Integer])
increment x = do num <- read x
write x (num + 1)
```

This function reads the value stored at the resource `x`

with `read`

,
increments it then writes the result back into the resource `x`

with
`write`

. We’ll see the types of `read`

and `write`

shortly
(see STrans Primitive operations). We can also create and delete resources:

```
makeAndIncrement : Integer -> STrans m Integer [] (const [])
makeAndIncrement init = do var <- new init
increment var
x <- read var
delete var
pure x
```

The type of `makeAndIncrement`

states that it has *no* resources available on
entry (`[]`

) or exit (`const []`

). It creates a new `State`

resource with
`new`

(which takes an initial value for the resource), increments the value,
reads it back, then deletes it using `delete`

, returning the final value
of the resource. Again, we’ll see the types of `new`

and `delete`

shortly.

The `m`

argument to `STrans`

(of type `Type -> Type`

) is the *computation context* in
which the function can be run. Here, the type level variable indicates that we
can run it in *any* context. We can run it in the identity context with
`runPure`

. For example, try entering the above definitions in a file
`Intro.idr`

then running the following at the REPL:

```
*Intro> runPure (makeAndIncrement 93)
94 : Integer
```

It’s a good idea to take an interactive, type-driven approach to implementing
`STrans`

programs. For example, after creating the resource with `new init`

,
you can leave a *hole* for the rest of the program to see how creating the
resource has affected the type:

```
makeAndIncrement : Integer -> STrans m Integer [] (const [])
makeAndIncrement init = do var <- new init
?whatNext
```

If you check the type of `?whatNext`

, you’ll see that there is now
a resource available, `var`

, and that by the end of the function there
should be no resource available:

```
init : Integer
m : Type -> Type
var : Var
--------------------------------------
whatNext : STrans m Integer [var ::: State Integer] (\value => [])
```

These small examples work in any computation context `m`

. However, usually,
we are working in a more restricted context. For example, we might want to
write programs which only work in a context that supports interactive
programs. For this, we’ll need to see how to *lift* operations from the
underlying context.

#### Lifting: Using the computation context¶

Let’s say that, instead of passing an initial integer to `makeAndIncrement`

,
we want to read it in from the console. Then, instead of working in a generic
context `m`

, we can work in the specific context `IO`

:

```
ioMakeAndIncrement : STrans IO () [] (const [])
```

This gives us access to `IO`

operations, via the `lift`

function. We
can define `ioMakeAndIncrement`

as follows:

```
ioMakeAndIncrement : STrans IO () [] (const [])
ioMakeAndIncrement
= do lift $ putStr "Enter a number: "
init <- lift $ getLine
var <- new (cast init)
lift $ putStrLn ("var = " ++ show !(read var))
increment var
lift $ putStrLn ("var = " ++ show !(read var))
delete var
```

The `lift`

function allows us to use funtions from the underlying
computation context (`IO`

here) directly. Again, we’ll see the exact type
of `lift`

shortly.

!-notation

In `ioMakeAndIncrement`

we’ve used `!(read var)`

to read from the
resource. You can read about this `!`

-notation in the main Idris tutorial
(see Monads and do-notation). In short, it allows us to use an `STrans`

function inline, rather than having to bind the result to a variable
first.

Conceptually, at least, you can think of it as having the following type:

```
(!) : STrans m a state_in state_out -> a
```

It is syntactic sugar for binding a variable immediately before the
current action in a `do`

block, then using that variable in place of
the `!`

-expression.

In general, though, it’s bad practice to use a *specific* context like
`IO`

. Firstly, it requires us to sprinkle `lift`

liberally throughout
our code, which hinders readability. Secondly, and more importantly, it will
limit the safety of our functions, as we’ll see in the next section
(State Machines in Types).

So, instead, we define *interfaces* to restrict the computation context.
For example, `Control.ST`

defines a `ConsoleIO`

interface which
provides the necessary methods for performing basic console interaction:

```
interface ConsoleIO (m : Type -> Type) where
putStr : String -> STrans m () res (const res)
getStr : STrans m String res (const res)
```

That is, we can write to and read from the console with any available
resources `res`

, and neither will affect the available resources.
This has the following implementation for `IO`

:

```
ConsoleIO IO where
putStr str = lift (Interactive.putStr str)
getStr = lift Interactive.getLine
```

Now, we can define `ioMakeAndIncrement`

as follows:

```
ioMakeAndIncrement : ConsoleIO io => STrans io () [] (const [])
ioMakeAndIncrement
= do putStr "Enter a number: "
init <- getStr
var <- new (cast init)
putStrLn ("var = " ++ show !(read var))
increment var
putStrLn ("var = " ++ show !(read var))
delete var
```

Instead of working in `IO`

specifically, this works in a generic context
`io`

, provided that there is an implementation of `ConsoleIO`

for that
context. This has several advantages over the first version:

- All of the calls to
`lift`

are in the implementation of the interface, rather than`ioMakeAndIncrement`

- We can provide alternative implementations of
`ConsoleIO`

, perhaps supporting exceptions or logging in addition to basic I/O. - As we’ll see in the next section (State Machines in Types), it will allow us to define safe APIs for manipulating specific resources more precisely.

Earlier, we used `runPure`

to run `makeAndIncrement`

in the identity
context. Here, we use `run`

, which allows us to execute an `STrans`

program
in any context (as long as it has an implementation of `Applicative`

) and we
can execute `ioMakeAndIncrement`

at the REPL as follows:

```
*Intro> :exec run ioMakeAndIncrement
Enter a number: 93
var = 93
var = 94
```

#### Manipulating `State`

with dependent types¶

In our first example of `State`

, when we incremented the value its
*type* remained the same. However, when we’re working with
*dependent* types, updating a state may also involve updating its type.
For example, if we’re adding an element to a vector stored in a state,
its length will change:

```
addElement : (vec : Var) -> (item : a) ->
STrans m () [vec ::: State (Vect n a)]
(const [vec ::: State (Vect (S n) a)])
addElement vec item = do xs <- read vec
write vec (item :: xs)
```

Note that you’ll need to `import Data.Vect`

to try this example.

Updating a state directly with `update`

Rather than using `read`

and `write`

separately, you can also
use `update`

which reads from a `State`

, applies a function to it,
then writes the result. Using `update`

you could write `addElement`

as follows:

```
addElement : (vec : Var) -> (item : a) ->
STrans m () [vec ::: State (Vect n a)]
(const [vec ::: State (Vect (S n) a)])
addElement vec item = update vec (item ::)
```

We don’t always know *how* exactly the type will change in the course of a
sequence actions, however. For example, if we have a state containing a
vector of integers, we might read an input from the console and only add it
to the vector if the input is a valid integer. Somehow, we need a different
type for the output state depending on whether reading the integer was
successful, so neither of the following types is quite right:

```
readAndAdd_OK : ConsoleIO io => (vec : Var) ->
STrans m () -- Returns an empty tuple
[vec ::: State (Vect n Integer)]
(const [vec ::: State (Vect (S n) Integer)])
readAndAdd_Fail : ConsoleIO io => (vec : Var) ->
STrans m () -- Returns an empty tuple
[vec ::: State (Vect n Integer)]
(const [vec ::: State (Vect n Integer)])
```

Remember, though, that the *output* resource types can be *computed* from
the result of a function. So far, we’ve used `const`

to note that the
output resources are always the same, but here, instead, we can use a type
level function to *calculate* the output resources. We start by returning
a `Bool`

instead of an empty tuple, which is `True`

if reading the input
was successful, and leave a *hole* for the output resources:

```
readAndAdd : ConsoleIO io => (vec : Var) ->
STrans m Bool [vec ::: State (Vect n Integer)]
?output_res
```

If you check the type of `?output_res`

, you’ll see that Idris expects
a function of type `Bool -> Resources`

, meaning that the output resource
type can be different depending on the result of `readAndAdd`

:

```
n : Nat
m : Type -> Type
io : Type -> Type
constraint : ConsoleIO io
vec : Var
--------------------------------------
output_res : Bool -> Resources
```

So, the output resource is either a `Vect n Integer`

if the input is
invalid (i.e. `readAndAdd`

returns `False`

) or a `Vect (S n) Integer`

if the input is valid. We can express this in the type as follows:

```
readAndAdd : ConsoleIO io => (vec : Var) ->
STrans io Bool [vec ::: State (Vect n Integer)]
(\res => [vec ::: State (if res then Vect (S n) Integer
else Vect n Integer)])
```

Then, when we implement `readAndAdd`

we need to return the appropriate
value for the output state. If we’ve added an item to the vector, we need to
return `True`

, otherwise we need to return `False`

:

```
readAndAdd : ConsoleIO io => (vec : Var) ->
STrans io Bool [vec ::: State (Vect n Integer)]
(\res => [vec ::: State (if res then Vect (S n) Integer
else Vect n Integer)])
readAndAdd vec = do putStr "Enter a number: "
num <- getStr
if all isDigit (unpack num)
then do
update vec ((cast num) ::)
pure True -- added an item, so return True
else pure False -- didn't add, so return False
```

There is a slight difficulty if we’re developing interactively, which is
that if we leave a hole, the required output state isn’t easily visible
until we know the value that’s being returned. For example. in the following
incomplete definition of `readAndAdd`

we’ve left a hole for the
successful case:

```
readAndAdd vec = do putStr "Enter a number: "
num <- getStr
if all isDigit (unpack num)
then ?whatNow
else pure False
```

We can look at the type of `?whatNow`

, but it is unfortunately rather less
than informative:

```
vec : Var
n : Nat
io : Type -> Type
constraint : ConsoleIO io
num : String
--------------------------------------
whatNow : STrans io Bool [vec ::: State (Vect (S n) Integer)]
(\res =>
[vec :::
State (ifThenElse res
(Delay (Vect (S n) Integer))
(Delay (Vect n Integer)))])
```

The problem is that we’ll only know the required output state when we know
the value we’re returning. To help with interactive development, `Control.ST`

provides a function `returning`

which allows us to specify the return
value up front, and to update the state accordingly. For example, we can
write an incomplete `readAndAdd`

as follows:

```
readAndAdd vec = do putStr "Enter a number: "
num <- getStr
if all isDigit (unpack num)
then returning True ?whatNow
else pure False
```

This states that, in the successful branch, we’ll be returning `True`

, and
`?whatNow`

should explain how to update the states appropriately so that
they are correct for a return value of `True`

. We can see this by checking
the type of `?whatNow`

, which is now a little more informative:

```
vec : Var
n : Nat
io : Type -> Type
constraint : ConsoleIO io
num : String
--------------------------------------
whatnow : STrans io () [vec ::: State (Vect n Integer)]
(\value => [vec ::: State (Vect (S n) Integer)])
```

This type now shows, in the output resource list of `STrans`

,
that we can complete the definition by adding an item to `vec`

, which
we can do as follows:

```
readAndAdd vec = do putStr "Enter a number: "
num <- getStr
if all isDigit (unpack num)
then returning True (update vec ((cast num) ::))
else returning False (pure ()) -- returning False, so no state update required
```

`STrans`

Primitive operations¶

Now that we’ve written a few small examples of `STrans`

functions, it’s
a good time to look more closely at the types of the state manipulation
functions we’ve used. First, to read and write states, we’ve used
`read`

and `write`

:

```
read : (lbl : Var) -> {auto prf : InState lbl (State ty) res} ->
STrans m ty res (const res)
write : (lbl : Var) -> {auto prf : InState lbl ty res} ->
(val : ty') ->
STrans m () res (const (updateRes res prf (State ty')))
```

These types may look a little daunting at first, particularly due to the
implicit `prf`

argument, which has the following type:

```
prf : InState lbl (State ty) res
```

This relies on a predicate `InState`

. A value of type `InState x ty res`

means that the reference `x`

must have type `ty`

in the list of
resources `res`

. So, in practice, all this type means is that we can
only read or write a resource if a reference to it exists in the list of
resources.

Given a resource label `res`

, and a proof that `res`

exists in a list
of resources, `updateRes`

will update the type of that resource. So,
the type of `write`

states that the type of the resource will be updated
to the type of the given value.

The type of `update`

is similar to that for `read`

and `write`

, requiring
that the resource has the input type of the given function, and updating it to
have the output type of the function:

```
update : (lbl : Var) -> {auto prf : InState lbl (State ty) res} ->
(ty -> ty') ->
STrans m () res (const (updateRes res prf (State ty')))
```

The type of `new`

states that it returns a `Var`

, and given an initial
value of type `state`

, the output resources contains a new resource
of type `State state`

:

```
new : (val : state) ->
STrans m Var res (\lbl => (lbl ::: State state) :: res)
```

It’s important that the new resource has type `State state`

, rather than
merely `state`

, because this will allow us to hide implementation details
of APIs. We’ll see more about what this means in the next section,
State Machines in Types.

The type of `delete`

states that the given label will be removed from
the list of resources, given an implicit proof that the label exists in
the input resources:

```
delete : (lbl : Var) -> {auto prf : InState lbl (State st) res} ->
STrans m () res (const (drop res prf))
```

Here, `drop`

is a type level function which updates the resource list,
removing the given resource `lbl`

from the list.

We’ve used `lift`

to run functions in the underlying context. It has the
following type:

```
lift : Monad m => m t -> STrans m t res (const res)
```

Given a `result`

value, `pure`

is an `STrans`

program which produces
that value, provided that the current list of resources is correct when
producing that value:

```
pure : (result : ty) -> STrans m ty (out_fn result) out_fn
```

We can use `returning`

to break down returning a value from an
`STrans`

functions into two parts: providing the value itself, and updating
the resource list so that it is appropriate for returning that value:

```
returning : (result : ty) ->
STrans m () res (const (out_fn result)) ->
STrans m ty res out_fn
```

Finally, we’ve used `run`

and `runPure`

to execute `STrans`

functions
in a specific context. `run`

will execute a function in any context,
provided that there is an `Applicative`

implementation for that context,
and `runPure`

will execute a function in the identity context:

```
run : Applicative m => STrans m a [] (const []) -> m a
runPure : STrans Basics.id a [] (const []) -> a
```

Note that in each case, the input and output resource list must be empty.
There’s no way to provide an initial resource list, or extract the final
resources. This is deliberate: it ensures that *all* resource management is
carried out in the controlled `STrans`

environment and, as we’ll see, this
allows us to implement safe APIs with precise types explaining exactly how
resources are tracked throughout a program.

These functions provide the core of the `ST`

library; there are some
others which we’ll encounter later, for more advanced situations, but the
functions we have seen so far already allow quite sophisticated state-aware
programming and reasoning in Idris.

#### ST: Representing state transitions directly¶

We’ve seen a few examples of small `STrans`

functions now, and
their types can become quite verbose given that we need to provide explicit
input and output resource lists. This is convenient for giving types for
the primitive operations, but for more general use it’s much more convenient
to be able to express *transitions* on individual resources, rather than
giving input and output resource lists in full. We can do this with
`ST`

:

```
ST : (m : Type -> Type) ->
(resultType : Type) ->
List (Action resultType) -> Type
```

`ST`

is a type level function which computes an appropriate `STrans`

type given a list of *actions*, which describe transitions on resources.
An `Action`

in a function type can take one of the following forms (plus
some others which we’ll see later in the tutorial):

`lbl ::: ty`

expresses that the resource`lbl`

begins and ends in the state`ty`

`lbl ::: ty_in :-> ty_out`

expresses that the resource`lbl`

begins in state`ty_in`

and ends in state`ty_out`

`lbl ::: ty_in :-> (\res -> ty_out)`

expresses that the resource`lbl`

begins in state`ty_in`

and ends in a state`ty_out`

, where`ty_out`

is computed from the result of the function`res`

.

So, we can write some of the function types we’ve seen so far as follows:

```
increment : (x : Var) -> ST m () [x ::: State Integer]
```

That is, `increment`

begins and ends with `x`

in state `State Integer`

.

```
makeAndIncrement : Integer -> ST m Integer []
```

That is, `makeAndIncrement`

begins and ends with no resources.

```
addElement : (vec : Var) -> (item : a) ->
ST m () [vec ::: State (Vect n a) :-> State (Vect (S n) a)]
```

That is, `addElement`

changes `vec`

from `State (Vect n a)`

to
`State (Vect (S n) a)`

.

```
readAndAdd : ConsoleIO io => (vec : Var) ->
ST io Bool
[vec ::: State (Vect n Integer) :->
\res => State (if res then Vect (S n) Integer
else Vect n Integer)]
```

By writing the types in this way, we express the minimum necessary to explain
how each function affects the overall resource state. If there is a resource
update depending on a result, as with `readAndAdd`

, then we need to describe
it in full. Otherwise, as with `increment`

and `makeAndIncrement`

, we can
write the input and output resource lists without repetition.

An `Action`

can also describe *adding* and *removing* states:

`add ty`

, assuming the operation returns a`Var`

, adds a new resource of type`ty`

.`remove lbl ty`

expresses that the operation removes the resource named`lbl`

, beginning in state`ty`

from the resource list.

So, for example, we can write:

```
newState : ST m Var [add (State Int)]
removeState : (lbl : Var) -> ST m () [remove lbl (State Int)]
```

The first of these, `newState`

, returns a new resource label, and adds that
resource to the list with type `State Int`

. The second, `removeState`

,
given a label `lbl`

, removes the resource from the list. These types are
equivalent to the following:

```
newState : STrans m Var [] (\lbl => [lbl ::: State Int])
removeState : (lbl : Var) -> STrans m () [lbl ::: State Int] (const [])
```

These are the primitive methods of constructing an `Action`

. Later, we will
encounter some other ways using type level functions to help with readability.

In the remainder of this tutorial, we will generally use `ST`

except on
the rare occasions we need the full precision of `STrans`

. In the next
section, we’ll see how to use the facilities provided by `ST`

to write
a precise API for a system with security properties: a data store requiring
a login.

### State Machines in Types¶

In the introduction, we saw the following state transition diagram representing the (abstract) states of a data store, and the actions we can perform on the store:

We say that these are the *abstract* states of the store, because the concrete
state will contain a lot more information: for example, it might contain
user names, hashed passwords, the store contents, and so on. However, as far
as we are concerned for the actions `login`

, `logout`

and `readSecret`

,
it’s whether we are logged in or not which affects which are valid.

We’ve seen how to manipulate states using `ST`

, and some small examples
of dependent types in states. In this section, we’ll see how to use
`ST`

to provide a safe API for the data store. In the API, we’ll encode
the above diagram in the types, in such a way that we can only execute the
operations `login`

, `logout`

and `readSecret`

when the state is
valid.

So far, we’ve used `State`

and the primitive operations, `new`

, `read`

,
`write`

and `delete`

to manipulate states. For the data store API,
however, we’ll begin by defining an *interface* (see Interfaces in
the Idris tutorial) which describes the operations on the store, and explains
in their types exactly when each operation is valid, and how it affects
the store’s state. By using an interface, we can be sure that
this is the *only* way to access the store.

#### Defining an interface for the data store¶

We’ll begin by defining a data type, in a file `Login.idr`

, which represents
the two abstract states of the store, either `LoggedOut`

or `LoggedIn`

:

```
data Access = LoggedOut | LoggedIn
```

We can define a data type for representing the current state of a store, holding all of the necessary information (this might be user names, hashed passwords, store contents and so on) and parameterise it by the logged in status of the store:

```
Store : Access -> Type
```

Rather than defining a concrete type now, however, we’ll include this in
a data store *interface* and define a concrete type later:

```
interface DataStore (m : Type -> Type) where
Store : Access -> Type
```

We can continue to populate this interface with operations on the store. Among
other advantages, by separating the *interface* from its *implementation* we
can provide different concrete implementations for different contexts.
Furthermore, we can write programs which work with a store without needing
to know any details of how the store is implemented.

We’ll need to be able to `connect`

to a store, and `disconnect`

when
we’re done. Add the following methods to the `DataStore`

interface:

```
connect : ST m Var [add (Store LoggedOut)]
disconnect : (store : Var) -> ST m () [remove store (Store LoggedOut)]
```

The type of `connect`

says that it returns a new resource which has the
initial type `Store LoggedOut`

. Conversely, `disconnect`

, given a
resource in the state `Store LoggedOut`

, removes that resource.
We can see more clearly what `connect`

does by trying the following
(incomplete) definition:

```
doConnect : DataStore m => ST m () []
doConnect = do st <- connect
?whatNow
```

Note that we’re working in a *generic* context `m`

, constrained so that
there must be an implementation of `DataStore`

for `m`

to be able to
execute `doConnect`

.
If we check the type of `?whatNow`

, we’ll see that the remaining
operations begin with a resource `st`

in the state `Store LoggedOut`

,
and we need to finish with no resources.

```
m : Type -> Type
constraint : DataStore m
st : Var
--------------------------------------
whatNow : STrans m () [st ::: Store LoggedOut] (\result => [])
```

Then, we can remove the resource using `disconnect`

:

```
doConnect : DataStore m => ST m () []
doConnect = do st <- connect
disconnect st
?whatNow
```

Now checking the type of `?whatNow`

shows that we have no resources
available:

```
m : Type -> Type
constraint : DataStore m
st : Var
--------------------------------------
whatNow : STrans m () [] (\result => [])
```

To continue our implementation of the `DataStore`

interface, next we’ll add a
method for reading the secret data. This requires that the `store`

is in the
state `Store LoggedIn`

:

```
readSecret : (store : Var) -> ST m String [store ::: Store LoggedIn]
```

At this point we can try writing a function which connects to a store,
reads the secret, then disconnects. However, it will be unsuccessful, because
`readSecret`

requires us to be logged in:

```
badGet : DataStore m => ST m () []
badGet = do st <- connect
secret <- readSecret st
disconnect st
```

This results in the following error, because `connect`

creates a new
store in the `LoggedOut`

state, and `readSecret`

requires the store
to be in the `LoggedIn`

state:

```
When checking an application of function Control.ST.>>=:
Error in state transition:
Operation has preconditions: [st ::: Store LoggedOut]
States here are: [st ::: Store LoggedIn]
Operation has postconditions: \result => []
Required result states here are: \result => []
```

The error message explains how the required input states (the preconditions)
and the required output states (the postconditions) differ from the states
in the operation. In order to use `readSecret`

, we’ll need a way to get
from a `Store LoggedOut`

to a `Store LoggedIn`

. As a first attempt,
we can try the following type for `login`

:

```
login : (store : Var) -> ST m () [store ::: Store LoggedOut :-> Store LoggedIn] -- Incorrect type!
```

Note that in the *interface* we say nothing about *how* `login`

works;
merely how it affects the overall state. Even so, there is a problem with
the type of `login`

, because it makes the assumption that it will always
succeed. If it fails - for example because the implementation prompts for
a password and the user enters the password incorrectly - then it must not
result in a `LoggedIn`

store.

Instead, therefore, `login`

will return whether logging in was successful,
via the following type;

```
data LoginResult = OK | BadPassword
```

Then, we can *calculate* the result state (see Manipulating State with dependent types) from the
result. Add the following method to the `DataStore`

interface:

```
login : (store : Var) ->
ST m LoginResult [store ::: Store LoggedOut :->
(\res => Store (case res of
OK => LoggedIn
BadPassword => LoggedOut))]
```

If `login`

was successful, then the state after `login`

is
`Store LoggedIn`

. Otherwise, the state is `Store LoggedOut`

.

To complete the interface, we’ll add a method for logging out of the store.
We’ll assume that logging out is always successful, and moves the store
from the `Store LoggedIn`

state to the `Store LoggedOut`

state.

```
logout : (store : Var) -> ST m () [store ::: Store LoggedIn :-> Store LoggedOut]
```

This completes the interface, repeated in full for reference below:

```
interface DataStore (m : Type -> Type) where
Store : Access -> Type
connect : ST m Var [add (Store LoggedOut)]
disconnect : (store : Var) -> ST m () [remove store (Store LoggedOut)]
readSecret : (store : Var) -> ST m String [store ::: Store LoggedIn]
login : (store : Var) ->
ST m LoginResult [store ::: Store LoggedOut :->
(\res => Store (case res of
OK => LoggedIn
BadPassword => LoggedOut))]
logout : (store : Var) -> ST m () [store ::: Store LoggedIn :-> Store LoggedOut]
```

Before we try creating any implementations of this interface, let’s see how we can write a function with it, to log into a data store, read the secret if login is successful, then log out again.

#### Writing a function with the data store¶

As an example of working with the `DataStore`

interface, we’ll write a
function `getData`

, which connects to a store in order to read some data from
it. We’ll write this function interactively, step by step, using the types of
the operations to guide its development. It has the following type:

```
getData : (ConsoleIO m, DataStore m) => ST m () []
```

This type means that there are no resources available on entry or exit.
That is, the overall list of actions is `[]`

, meaning that at least
externally, the function has no overall effect on the resources. In other
words, for every resource we create during `getData`

, we’ll also need to
delete it before exit.

Since we want to use methods of the `DataStore`

interface, we’ll
constraint the computation context `m`

so that there must be an
implementation of `DataStore`

. We also have a constraint `ConsoleIO m`

so that we can display any data we read from the store, or any error
messages.

We start by connecting to the store, creating a new resource `st`

, then
trying to `login`

:

```
getData : (ConsoleIO m, DataStore m) => ST m () []
getData = do st <- connect
ok <- login st
?whatNow
```

Logging in will either succeed or fail, as reflected by the value of
`ok`

. If we check the type of `?whatNow`

, we’ll see what state the
store currently has:

```
m : Type -> Type
constraint : ConsoleIO m
constraint1 : DataStore m
st : Var
ok : LoginResult
--------------------------------------
whatNow : STrans m () [st ::: Store (case ok of
OK => LoggedIn
BadPassword => LoggedOut)]
(\result => [])
```

The current state of `st`

therefore depends on the value of `ok`

,
meaning that we can make progress by case splitting on `ok`

:

```
getData : (ConsoleIO m, DataStore m) => ST m () []
getData = do st <- connect
ok <- login st
case ok of
OK => ?whatNow_1
BadPassword => ?whatNow_2
```

The types of the holes in each branch, `?whatNow_1`

and `?whatNow_2`

,
show how the state changes depending on whether logging in was successful.
If it succeeded, the store is `LoggedIn`

:

```
--------------------------------------
whatNow_1 : STrans m () [st ::: Store LoggedIn] (\result => [])
```

On the other hand, if it failed, the store is `LoggedOut`

:

```
--------------------------------------
whatNow_2 : STrans m () [st ::: Store LoggedOut] (\result => [])
```

In `?whatNow_1`

, since we’ve successfully logged in, we can now read
the secret and display it to the console:

```
getData : (ConsoleIO m, DataStore m) => ST m () []
getData = do st <- connect
ok <- login st
case ok of
OK => do secret <- readSecret st
putStrLn ("Secret is: " ++ show secret)
?whatNow_1
BadPassword => ?whatNow_2
```

We need to finish the `OK`

branch with no resources available. We can
do this by logging out of the store then disconnecting:

```
getData : (ConsoleIO m, DataStore m) => ST m () []
getData = do st <- connect
ok <- login st
case ok of
OK => do secret <- readSecret st
putStrLn ("Secret is: " ++ show secret)
logout st
disconnect st
BadPassword => ?whatNow_2
```

Note that we *must* `logout`

of `st`

before calling `disconnect`

,
because `disconnect`

requires that the store is in the `LoggedOut`

state.

Furthermore, we can’t simply use `delete`

to remove the resource, as
we did with the `State`

examples in the previous section, because
`delete`

only works when the resource has type `State ty`

, for some
type `ty`

. If we try to use `delete`

instead of `disconnect`

, we’ll
see an error message like the following:

```
When checking argument prf to function Control.ST.delete:
Can't find a value of type
InState st (State st) [st ::: Store LoggedOut]
```

In other words, the type checker can’t find a proof that the resource
`st`

has a type of the form `State st`

, because its type is
`Store LoggedOut`

. Since `Store`

is part of the `DataStore`

interface,
we *can’t* yet know the concrete representation of the `Store`

, so we
need to remove the resource via the interface, with `disconnect`

, rather
than directly with `delete`

.

We can complete `getData`

as follows, using a pattern matching bind
alternative (see the Idris tutorial, Monads and do-notation) rather than a
`case`

statement to catch the possibility of an error with `login`

:

```
getData : (ConsoleIO m, DataStore m) => ST m () []
getData = do st <- connect
OK <- login st
| BadPassword => do putStrLn "Failure"
disconnect st
secret <- readSecret st
putStrLn ("Secret is: " ++ show secret)
logout st
disconnect st
```

We can’t yet try this out, however, because we don’t have any implementations
of `DataStore`

! If we try to execute it in an `IO`

context, for example,
we’ll get an error saying that there’s no implementation of `DataStore IO`

:

```
*Login> :exec run {m = IO} getData
When checking an application of function Control.ST.run:
Can't find implementation for DataStore IO
```

The final step in implementing a data store which correctly follows the
state transition diagram, therefore, is to provide an implementation
of `DataStore`

.

#### Implementing the interface¶

To execute `getData`

in `IO`

, we’ll need to provide an implementation
of `DataStore`

which works in the `IO`

context. We can begin as
follows:

```
implementation DataStore IO where
```

Then, we can ask Idris to populate the interface with skeleton definitions
for the necessary methods (press `Ctrl-Alt-A`

in Atom for “add definition”
or the corresponding shortcut for this in the Idris mode in your favourite
editor):

```
implementation DataStore IO where
Store x = ?DataStore_rhs_1
connect = ?DataStore_rhs_2
disconnect store = ?DataStore_rhs_3
readSecret store = ?DataStore_rhs_4
login store = ?DataStore_rhs_5
logout store = ?DataStore_rhs_6
```

The first decision we’ll need to make is how to represent the data store.
We’ll keep this simple, and store the data as a single `String`

, using
a hard coded password to gain access. So, we can define `Store`

as
follows, using a `String`

to represent the data no matter whether we
are `LoggedOut`

or `LoggedIn`

:

```
Store x = State String
```

Now that we’ve given a concrete type for `Store`

, we can implement operations
for connecting, disconnecting, and accessing the data. And, since we used
`State`

, we can use `new`

, `delete`

, `read`

and `write`

to
manipulate the store.

Looking at the types of the holes tells us how we need to manipulate the
state. For example, the `?DataStore_rhs_2`

hole tells us what we need
to do to implement `connect`

. We need to return a new `Var`

which
represents a resource of type `State String`

:

```
--------------------------------------
DataStore_rhs_2 : STrans IO Var [] (\result => [result ::: State String])
```

We can implement this by creating a new variable with some data for the
content of the store (we can use any `String`

for this) and returning
that variable:

```
connect = do store <- new "Secret Data"
pure store
```

For `disconnect`

, we only need to delete the resource:

```
disconnect store = delete store
```

For `readSecret`

, we need to read the secret data and return the
`String`

. Since we now know the concrete representation of the data is
a `State String`

, we can use `read`

to access the data directly:

```
readSecret store = read store
```

We’ll do `logout`

next and return to `login`

. Checking the hole
reveals the following:

```
store : Var
--------------------------------------
DataStore_rhs_6 : STrans IO () [store ::: State String] (\result => [store ::: State String])
```

So, in this minimal implementation, we don’t actually have to do anything!

```
logout store = pure ()
```

For `login`

, we need to return whether logging in was successful. We’ll
do this by prompting for a password, and returning `OK`

if it matches
a hard coded password, or `BadPassword`

otherwise:

```
login store = do putStr "Enter password: "
p <- getStr
if p == "Mornington Crescent"
then pure OK
else pure BadPassword
```

For reference, here is the complete implementation which allows us to
execute a `DataStore`

program at the REPL:

```
implementation DataStore IO where
Store x = State String
connect = do store <- new "Secret Data"
pure store
disconnect store = delete store
readSecret store = read store
login store = do putStr "Enter password: "
p <- getStr
if p == "Mornington Crescent"
then pure OK
else pure BadPassword
logout store = pure ()
```

Finally, we can try this at the REPL as follows (Idris defaults to the
`IO`

context at the REPL if there is an implementation available, so no
need to give the `m`

argument explicitly here):

```
*Login> :exec run getData
Enter password: Mornington Crescent
Secret is: "Secret Data"
*Login> :exec run getData
Enter password: Dollis Hill
Failure
```

We can only use `read`

, `write`

, `new`

and `delete`

on a resource
with a `State`

type. So, *within* the implementation of `DataStore`

,
or anywhere where we know the context is `IO`

, we can access the data store
however we like: this is where the internal details of `DataStore`

are
implemented. However, if we merely have a constraint `DataStore m`

, we can’t
know how the store is implemented, so we can only access via the API given
by the `DataStore`

interface.

It is therefore good practice to use a *generic* context `m`

for functions
like `getData`

, and constrain by only the interfaces we need, rather than
using a concrete context `IO`

.

We’ve now seen how to manipulate states, and how to encapsulate state
transitions for a specific system like the data store in an interface.
However, realistic systems will need to *compose* state machines. We’ll
either need to use more than one state machine at a time, or implement one
state machine in terms of one or more others. We’ll see how to achieve this
in the next section.

### Composing State Machines¶

In the previous section, we defined a `DataStore`

interface and used it
to implement the following small program which allows a user to log in to
the store then display the store’s contents;

```
getData : (ConsoleIO m, DataStore m) => ST m () []
getData = do st <- connect
OK <- login st
| BadPassword => do putStrLn "Failure"
disconnect st
secret <- readSecret st
putStrLn ("Secret is: " ++ show secret)
logout st
disconnect st
```

This function only uses one state, the store itself. Usually, though, larger programs have lots of states, and might add, delete and update states over the course of its execution. Here, for example, a useful extension might be to loop forever, keeping count of the number of times there was a login failure in a state.

Furthermore, we may have *hierarchies* of state machines, in that one
state machine could be implemented by composing several others. For
example, we can have a state machine representing the state of a
graphics system, and use this to implement a *higher level* graphics API
such as turtle graphics, which uses the graphics system plus some additional
state for the turtle.

In this section, we’ll see how to work with multiple states, and how to
compose state machines to make higher level state machines. We’ll begin by
seeing how to add a login failure counter to `getData`

.

#### Working with multiple resources¶

To see how to work with multiple resources, we’ll modify `getData`

so
that it loops, and counts the total number of times the user fails to
log in. For example, if we write a `main`

program which initialises the
count to zero, a session might run as follows:

```
*LoginCount> :exec main
Enter password: Mornington Crescent
Secret is: "Secret Data"
Enter password: Dollis Hill
Failure
Number of failures: 1
Enter password: Mornington Crescent
Secret is: "Secret Data"
Enter password: Codfanglers
Failure
Number of failures: 2
...
```

We’ll start by adding a state resource to `getData`

to keep track of the
number of failures:

```
getData : (ConsoleIO m, DataStore m) =>
(failcount : Var) -> ST m () [failcount ::: State Integer]
```

Type checking `getData`

If you’re following along in the code, you’ll find that `getData`

no longer compiles when you update this type. That is to be expected!
For the moment, comment out the definition of `getData`

. We’ll come back
to it shortly.

Then, we can create a `main`

program which initialises the state to `0`

and invokes `getData`

, as follows:

```
main : IO ()
main = run (do fc <- new 0
getData fc
delete fc)
```

We’ll start our implementation of `getData`

just by adding the new
argument for the failure count:

```
getData : (ConsoleIO m, DataStore m) =>
(failcount : Var) -> ST m () [failcount ::: State Integer]
getData failcount
= do st <- connect
OK <- login st
| BadPassword => do putStrLn "Failure"
disconnect st
secret <- readSecret st
putStrLn ("Secret is: " ++ show secret)
logout st
disconnect st
```

Unfortunately, this doesn’t type check, because we have the wrong resources
for calling `connect`

. The error messages shows how the resources don’t
match:

```
When checking an application of function Control.ST.>>=:
Error in state transition:
Operation has preconditions: []
States here are: [failcount ::: State Integer]
Operation has postconditions: \result => [result ::: Store LoggedOut] ++ []
Required result states here are: st2_fn
```

In other words, `connect`

requires that there are *no* resources on
entry, but we have *one*, the failure count!
This shouldn’t be a problem, though: the required resources are a *subset* of
the resources we have, after all, and the additional resources (here, the
failure count) are not relevant to `connect`

. What we need, therefore,
is a way to temporarily *hide* the additional resource.

We can achieve this with the `call`

function:

```
getData : (ConsoleIO m, DataStore m) =>
(failcount : Var) -> ST m () [failcount ::: State Integer]
getData failcount
= do st <- call connect
?whatNow
```

Here we’ve left a hole for the rest of `getData`

so that you can see the
effect of `call`

. It has removed the unnecessary parts of the resource
list for calling `connect`

, then reinstated them on return. The type of
`whatNow`

therefore shows that we’ve added a new resource `st`

, and still
have `failcount`

available:

```
failcount : Var
m : Type -> Type
constraint : ConsoleIO m
constraint1 : DataStore m
st : Var
--------------------------------------
whatNow : STrans m () [failcount ::: State Integer, st ::: Store LoggedOut]
(\result => [failcount ::: State Integer])
```

By the end of the function, `whatNow`

says that we need to have finished with
`st`

, but still have `failcount`

available. We can complete `getData`

so that it works with an additional state resource by adding `call`

whenever
we invoke one of the operations on the data store, to reduce the list of
resources:

```
getData : (ConsoleIO m, DataStore m) =>
(failcount : Var) -> ST m () [failcount ::: State Integer]
getData failcount
= do st <- call connect
OK <- call $ login st
| BadPassword => do putStrLn "Failure"
call $ disconnect st
secret <- call $ readSecret st
putStrLn ("Secret is: " ++ show secret)
call $ logout st
call $ disconnect st
```

This is a little noisy, and in fact we can remove the need for it by
making `call`

implicit. By default, you need to add the `call`

explicitly,
but if you import `Control.ST.ImplicitCall`

, Idris will insert `call`

where it is necessary.

```
import Control.ST.ImplicitCall
```

It’s now possible to write `getData`

exactly as before:

```
getData : (ConsoleIO m, DataStore m) =>
(failcount : Var) -> ST m () [failcount ::: State Integer]
getData failcount
= do st <- connect
OK <- login st
| BadPassword => do putStrLn "Failure"
disconnect st
secret <- readSecret st
putStrLn ("Secret is: " ++ show secret)
logout st
disconnect st
```

There is a trade off here: if you import `Control.ST.ImplicitCall`

then
functions which use multiple resources are much easier to read, because the
noise of `call`

has gone. On the other hand, Idris has to work a little
harder to type check your functions, and as a result it can take slightly
longer, and the error messages can be less helpful.

It is instructive to see the type of `call`

:

```
call : STrans m t sub new_f -> {auto res_prf : SubRes sub old} ->
STrans m t old (\res => updateWith (new_f res) old res_prf)
```

The function being called has a list of resources `sub`

, and
there is an implicit proof, `SubRes sub old`

that the resource list in
the function being called is a subset of the overall resource list. The
ordering of resources is allowed to change, although resources which
appear in `old`

can’t appear in the `sub`

list more than once (you will
get a type error if you try this).

The function `updateWith`

takes the *output* resources of the
called function, and updates them in the current resource list. It makes
an effort to preserve ordering as far as possible, although this isn’t
always possible if the called function does some complicated resource
manipulation.

Newly created resources in called functions

If the called function creates any new resources, these will typically
appear at the *end* of the resource list, due to the way `updateWith`

works. You can see this in the type of `whatNow`

in our incomplete
definition of `getData`

above.

Finally, we can update `getData`

so that it loops, and keeps
`failCount`

updated as necessary:

```
getData : (ConsoleIO m, DataStore m) =>
(failcount : Var) -> ST m () [failcount ::: State Integer]
getData failcount
= do st <- call connect
OK <- login st
| BadPassword => do putStrLn "Failure"
fc <- read failcount
write failcount (fc + 1)
putStrLn ("Number of failures: " ++ show (fc + 1))
disconnect st
getData failcount
secret <- readSecret st
putStrLn ("Secret is: " ++ show secret)
logout st
disconnect st
getData failcount
```

Note that here, we’re connecting and disconnecting on every iteration.
Another way to implement this would be to `connect`

first, then call
`getData`

, and implement `getData`

as follows:

```
getData : (ConsoleIO m, DataStore m) =>
(st, failcount : Var) -> ST m () [st ::: Store {m} LoggedOut, failcount ::: State Integer]
getData st failcount
= do OK <- login st
| BadPassword => do putStrLn "Failure"
fc <- read failcount
write failcount (fc + 1)
putStrLn ("Number of failures: " ++ show (fc + 1))
getData st failcount
secret <- readSecret st
putStrLn ("Secret is: " ++ show secret)
logout st
getData st failcount
```

It is important to add the explicit `{m}`

in the type of ```
Store {m}
LoggedOut
```

for `st`

, because this gives Idris enough information to know
which implementation of `DataStore`

to use to find the appropriate
implementation for `Store`

. Otherwise, if we only write `Store LoggedOut`

,
there’s no way to know that the `Store`

is linked with the computation
context `m`

.

We can then `connect`

and `disconnect`

only once, in `main`

:

```
main : IO ()
main = run (do fc <- new 0
st <- connect
getData st fc
disconnect st
delete fc)
```

By using `call`

, and importing `Control.ST.ImplicitCall`

, we can
write programs which use multiple resources, and reduce the list of
resources as necessary when calling functions which only use a subset of
the overall resources.

#### Composite resources: Hierarchies of state machines¶

We’ve now seen how to use multiple resources in one function, which is necessary for any realistic program which manipulates state. We can think of this as “horizontal” composition: using multiple resources at once. We’ll often also need “vertical” composition: implementing one resource in terms of one or more other resources.

We’ll see an example of this in this section. First, we’ll implement a
small API for graphics, in an interface `Draw`

, supporting:

- Opening a window, creating a double-buffered surface to draw on
- Drawing lines and rectangles onto a surface
- “Flipping” buffers, displaying the surface we’ve just drawn onto in the window
- Closing a window

Then, we’ll use this API to implement a higher level API for turtle graphics,
in an `interface`

.
This will require not only the `Draw`

interface, but also a representation
of the turtle state (location, direction and pen colour).

SDL bindings

For the examples in this section, you’ll need to install the (very basic!) SDL bindings for Idris, available from https://github.com/edwinb/SDL-idris. These bindings implement a small subset of the SDL API, and are for illustrative purposes only. Nevertheless, they are enough to implement small graphical programs and demonstrate the concepts of this section.

Once you’ve installed this package, you can start Idris with the
`-p sdl`

flag, for the SDL bindings, and the `-p contrib`

flag,
for the `Control.ST`

library.

##### The `Draw`

interface¶

We’re going to use the Idris SDL bindings for this API, so you’ll need
to import `Graphics.SDL`

once you’ve installed the bindings.
We’ll start by defining the `Draw`

interface, which includes a data type
representing a surface on which we’ll draw lines and rectangles:

```
interface Draw (m : Type -> Type) where
Surface : Type
```

We’ll need to be able to create a new `Surface`

by opening a window:

```
initWindow : Int -> Int -> ST m Var [add Surface]
```

However, this isn’t quite right. It’s possible that opening a window
will fail, for example if our program is running in a terminal without
a windowing system available. So, somehow, `initWindow`

needs to cope
with the possibility of failure. We can do this by returning a
`Maybe Var`

, rather than a `Var`

, and only adding the `Surface`

on success:

```
initWindow : Int -> Int -> ST m (Maybe Var) [addIfJust Surface]
```

This uses a type level function `addIfJust`

, defined in `Control.ST`

which returns an `Action`

that only adds a resource if the operation
succeeds (that is, returns a result of the form `Just val`

.

`addIfJust`

and `addIfRight`

`Control.ST`

defines functions for constructing new resources if an
operation succeeds. As well as `addIfJust`

, which adds a resource if
an operation returns `Just ty`

, there’s also `addIfRight`

:

```
addIfJust : Type -> Action (Maybe Var)
addIfRight : Type -> Action (Either a Var)
```

Each of these is implemented in terms of the following primitive action
`Add`

, which takes a function to construct a resource list from the result
of an operation:

```
Add : (ty -> Resources) -> Action ty
```

Using this, you can create your own actions to add resources
based on the result of an operation, if required. For example,
`addIfJust`

is implemented as follows:

```
addIfJust : Type -> Action (Maybe Var)
addIfJust ty = Add (maybe [] (\var => [var ::: ty]))
```

If we create windows, we’ll also need to be able to delete them:

```
closeWindow : (win : Var) -> ST m () [remove win Surface]
```

We’ll also need to respond to events such as keypresses and mouse clicks.
The `Graphics.SDL`

library provides an `Event`

type for this, and
we can `poll`

for events which returns the last event which occurred,
if any:

```
poll : ST m (Maybe Event) []
```

The remaining methods of `Draw`

are `flip`

, which flips the buffers
displaying everything that we’ve drawn since the previous `flip`

, and
two methods for drawing: `filledRectangle`

and `drawLine`

.

```
flip : (win : Var) -> ST m () [win ::: Surface]
filledRectangle : (win : Var) -> (Int, Int) -> (Int, Int) -> Col -> ST m () [win ::: Surface]
drawLine : (win : Var) -> (Int, Int) -> (Int, Int) -> Col -> ST m () [win ::: Surface]
```

We define colours as follows, as four components (red, green, blue, alpha):

```
data Col = MkCol Int Int Int Int
black : Col
black = MkCol 0 0 0 255
red : Col
red = MkCol 255 0 0 255
green : Col
green = MkCol 0 255 0 255
-- Also blue, yellow, magenta, cyan, white, similarly...
```

If you import `Graphics.SDL`

, you can implement the `Draw`

interface
using the SDL bindings as follows:

```
implementation Draw IO where
Surface = State SDLSurface
initWindow x y = do Just srf <- lift (startSDL x y)
| pure Nothing
var <- new srf
pure (Just var)
closeWindow win = do lift endSDL
delete win
flip win = do srf <- read win
lift (flipBuffers srf)
poll = lift pollEvent
filledRectangle win (x, y) (ex, ey) (MkCol r g b a)
= do srf <- read win
lift $ filledRect srf x y ex ey r g b a
drawLine win (x, y) (ex, ey) (MkCol r g b a)
= do srf <- read win
lift $ drawLine srf x y ex ey r g b a
```

In this implementation, we’ve used `startSDL`

to initialise a window, which,
returns `Nothing`

if it fails. Since the type of `initWindow`

states that
it adds a resource when it returns a value of the form `Just val`

, we
add the surface returned by `startSDL`

on success, and nothing on
failure. We can only successfully initialise if `startDSL`

succeeds.

Now that we have an implementation of `Draw`

, we can try writing some
functions for drawing into a window and execute them via the SDL bindings.
For example, assuming we have a surface `win`

to draw onto, we can write a
`render`

function as follows which draws a line onto a black background:

```
render : Draw m => (win : Var) -> ST m () [win ::: Surface {m}]
render win = do filledRectangle win (0,0) (640,480) black
drawLine win (100,100) (200,200) red
flip win
```

The `flip win`

at the end is necessary because the drawing primitives
are double buffered, to prevent flicker. We draw onto one buffer, off-screen,
and display the other. When we call `flip`

, it displays the off-screen
buffer, and creates a new off-screen buffer for drawing the next frame.

To include this in a program, we’ll write a main loop which renders our image and waits for an event to indicate the user wants to close the application:

```
loop : Draw m => (win : Var) -> ST m () [win ::: Surface {m}]
loop win = do render win
Just AppQuit <- poll
| _ => loop win
pure ()
```

Finally, we can create a main program which initialises a window, if possible, then runs the main loop:

```
drawMain : (ConsoleIO m, Draw m) => ST m () []
drawMain = do Just win <- initWindow 640 480
| Nothing => putStrLn "Can't open window"
loop win
closeWindow win
```

We can try this at the REPL using `run`

:

```
*Draw> :exec run drawMain
```

##### A higher level interface: `TurtleGraphics`

¶

Turtle graphics involves a “turtle” moving around the screen, drawing a line as it moves with a “pen”. A turtle has attributes describing its location, the direction it’s facing, and the current pen colour. There are commands for moving the turtle forwards, turning through an angle, and changing the pen colour, among other things. One possible interface would be the following:

```
interface TurtleGraphics (m : Type -> Type) where
Turtle : Type
start : Int -> Int -> ST m (Maybe Var) [addIfJust Turtle]
end : (t : Var) -> ST m () [Remove t Turtle]
fd : (t : Var) -> Int -> ST m () [t ::: Turtle]
rt : (t : Var) -> Int -> ST m () [t ::: Turtle]
penup : (t : Var) -> ST m () [t ::: Turtle]
pendown : (t : Var) -> ST m () [t ::: Turtle]
col : (t : Var) -> Col -> ST m () [t ::: Turtle]
render : (t : Var) -> ST m () [t ::: Turtle]
```

Like `Draw`

, we have a command for initialising the turtle (here called
`start`

) which might fail if it can’t create a surface for the turtle to
draw on. There is also a `render`

method, which is intended to render the
picture drawn so far in a window. One possible program with this interface
is the following, with draws a colourful square:

```
turtle : (ConsoleIO m, TurtleGraphics m) => ST m () []
turtle = with ST do
Just t <- start 640 480
| Nothing => putStr "Can't make turtle\n"
col t yellow
fd t 100; rt t 90
col t green
fd t 100; rt t 90
col t red
fd t 100; rt t 90
col t blue
fd t 100; rt t 90
render t
end t
```

`with ST do`

The purpose of `with ST do`

in `turtle`

is to disambiguate `(>>=)`

,
which could be either the version from the `Monad`

interface, or the
version from `ST`

. Idris can work this out itself, but it takes time to
try all of the possibilities, so the `with`

clause can
speed up type checking.

To implement the interface, we could try using `Surface`

to represent
the surface for the turtle to draw on:

```
implementation Draw m => TurtleGraphics m where
Turtle = Surface {m}
```

Knowing that a `Turtle`

is represented as a `Surface`

, we can use the
methods provided by `Draw`

to implement the turtle. Unfortunately, though,
this isn’t quite enough. We need to store more information: in particular, the
turtle has several attributes which we need to store somewhere.
So, not only do we need to represent the turtle as a `Surface`

, we need
to store some additional state. We can achieve this using a *composite*
resource.

##### Introducing composite resources¶

A *composite* resource is built up from a list of other resources, and
is implemented using the following type, defined by `Control.ST`

:

```
data Composite : List Type -> Type
```

If we have a composite resource, we can split it into its constituent
resources, and create new variables for each of those resources, using
the *split* function. For example:

```
splitComp : (comp : Var) -> ST m () [comp ::: Composite [State Int, State String]]
splitComp comp = do [int, str] <- split comp
?whatNow
```

The call `split comp`

extracts the `State Int`

and `State String`

from
the composite resource `comp`

, and stores them in the variables `int`

and `str`

respectively. If we check the type of `whatNow`

, we’ll see
how this has affected the resource list:

```
int : Var
str : Var
comp : Var
m : Type -> Type
--------------------------------------
whatNow : STrans m () [int ::: State Int, str ::: State String, comp ::: State ()]
(\result => [comp ::: Composite [State Int, State String]])
```

So, we have two new resources `int`

and `str`

, and the type of
`comp`

has been updated to the unit type, so currently holds no data.
This is to be expected: we’ve just extracted the data into individual
resources after all.

Now that we’ve extracted the individual resources, we can manipulate them
directly (say, incrementing the `Int`

and adding a newline to the
`String`

) then rebuild the composite resource using `combine`

:

```
splitComp : (comp : Var) ->
ST m () [comp ::: Composite [State Int, State String]]
splitComp comp = do [int, str] <- split comp
update int (+ 1)
update str (++ "\n")
combine comp [int, str]
?whatNow
```

As ever, we can check the type of `whatNow`

to see the effect of
`combine`

:

```
comp : Var
int : Var
str : Var
m : Type -> Type
--------------------------------------
whatNow : STrans m () [comp ::: Composite [State Int, State String]]
(\result => [comp ::: Composite [State Int, State String]])
```

The effect of `combine`

, therefore, is to take existing
resources and merge them into one composite resource. Before we run
`combine`

, the target resource must exist (`comp`

here) and must be
of type `State ()`

.

It is instructive to look at the types of `split`

and `combine`

to see
the requirements on resource lists they work with. The type of `split`

is the following:

```
split : (lbl : Var) -> {auto prf : InState lbl (Composite vars) res} ->
STrans m (VarList vars) res (\vs => mkRes vs ++ updateRes res prf (State ()))
```

The implicit `prf`

argument says that the `lbl`

being split must be
a composite resource. It returns a variable list, built from the composite
resource, and the `mkRes`

function makes a list of resources of the
appropriate types. Finally, `updateRes`

updates the composite resource to
have the type `State ()`

.

The `combine`

function does the inverse:

```
combine : (comp : Var) -> (vs : List Var) ->
{auto prf : InState comp (State ()) res} ->
{auto var_prf : VarsIn (comp :: vs) res} ->
STrans m () res (const (combineVarsIn res var_prf))
```

The implicit `prf`

argument here ensures that the target resource `comp`

has type `State ()`

. That is, we’re not overwriting any other data.
The implicit `var_prf`

argument is similar to `SubRes`

in `call`

, and
ensures that every variable we’re using to build the composite resource
really does exist in the current resource list.

We can use composite resources to implement our higher level `TurtleGraphics`

API in terms of `Draw`

, and any additional resources we need.

##### Implementing `Turtle`

¶

Now that we’ve seen how to build a new resource from an existing collection,
we can implement `Turtle`

using a composite resource, containing the
`Surface`

to draw on, and individual states for the pen colour and the
pen location and direction. We also have a list of lines, which describes
what we’ll draw onto the `Surface`

when we call `render`

:

```
Turtle = Composite [Surface {m}, -- surface to draw on
State Col, -- pen colour
State (Int, Int, Int, Bool), -- pen location/direction/d
State (List Line)] -- lines to draw on render
```

A `Line`

is defined as a start location, and end location, and a colour:

```
Line : Type
Line = ((Int, Int), (Int, Int), Col)
```

To implement `start`

, which creates a new `Turtle`

(or returns `Nothing`

if this is impossible), we begin by initialising the drawing surface then
all of the components of the state. Finally, we combine all of these
into a composite resource for the turtle:

```
start x y = do Just srf <- initWindow x y
| Nothing => pure Nothing
col <- new white
pos <- new (320, 200, 0, True)
lines <- new []
turtle <- new ()
combine turtle [srf, col, pos, lines]
pure (Just turtle)
```

To implement `end`

, which needs to dispose of the turtle,
we deconstruct the composite resource, close the window,
then remove each individual resource. Remember that we can only `delete`

a `State`

, so we need to `split`

the composite resource, close the
drawing surface cleanly with `closeWindow`

, then `delete`

the states:

```
end t = do [srf, col, pos, lines] <- split t
closeWindow srf; delete col; delete pos; delete lines; delete t
```

For the other methods, we need to `split`

the resource to get each
component, and `combine`

into a composite resource when we’re done.
As an example, here’s `penup`

:

```
penup t = do [srf, col, pos, lines] <- split t -- Split the composite resource
(x, y, d, _) <- read pos -- Deconstruct the pen position
write pos (x, y, d, False) -- Set the pen down flag to False
combine t [srf, col, pos, lines] -- Recombine the components
```

The remaining operations on the turtle follow a similar pattern. See
`samples/ST/Graphics/Turtle.idr`

in the Idris distribution for the full
details. It remains to render the image created by the turtle:

```
render t = do [srf, col, pos, lines] <- split t -- Split the composite resource
filledRectangle srf (0, 0) (640, 480) black -- Draw a background
drawAll srf !(read lines) -- Render the lines drawn by the turtle
flip srf -- Flip the buffers to display the image
combine t [srf, col, pos, lines]
Just ev <- poll
| Nothing => render t -- Keep going until a key is pressed
case ev of
KeyUp _ => pure () -- Key pressed, so quit
_ => render t
where drawAll : (srf : Var) -> List Line -> ST m () [srf ::: Surface {m}]
drawAll srf [] = pure ()
drawAll srf ((start, end, col) :: xs)
= do drawLine srf start end col -- Draw a line in the appropriate colour
drawAll srf xs
```

### Example: Network Socket Programming¶

The POSIX sockets API supports communication between processes across a
network. A *socket* represents an endpoint of a network communication, and can be
in one of several states:

`Ready`

, the initial state`Bound`

, meaning that it has been bound to an address ready for incoming connections`Listening`

, meaning that it is listening for incoming connections`Open`

, meaning that it is ready for sending and receiving data;`Closed`

, meaning that it is no longer active.

The following diagram shows how the operations provided by the API modify the
state, where `Ready`

is the initial state:

If a connection is `Open`

, then we can also `send`

messages to the
other end of the connection, and `recv`

messages from it.

The `contrib`

package provides a module `Network.Socket`

which
provides primitives for creating sockets and sending and receiving
messages. It includes the following functions:

```
bind : (sock : Socket) -> (addr : Maybe SocketAddress) -> (port : Port) -> IO Int
connect : (sock : Socket) -> (addr : SocketAddress) -> (port : Port) -> IO ResultCode
listen : (sock : Socket) -> IO Int
accept : (sock : Socket) -> IO (Either SocketError (Socket, SocketAddress))
send : (sock : Socket) -> (msg : String) -> IO (Either SocketError ResultCode)
recv : (sock : Socket) -> (len : ByteLength) -> IO (Either SocketError (String, ResultCode))
close : Socket -> IO ()
```

These functions cover the state transitions in the diagram above, but none of them explain how the operations affect the state! It’s perfectly possible, for example, to try to send a message on a socket which is not yet ready, or to try to receive a message after the socket is closed.

Using `ST`

, we can provide a better API which explains exactly how
each operation affects the state of a connection. In this section, we’ll
define a sockets API, then use it to implement an “echo” server which
responds to requests from a client by echoing back a single message sent
by the client.

#### Defining a `Sockets`

interface¶

Rather than using `IO`

for low level socket programming, we’ll implement
an interface using `ST`

which describes precisely how each operation
affects the states of sockets, and describes when sockets are created
and removed. We’ll begin by creating a type to describe the abstract state
of a socket:

```
data SocketState = Ready | Bound | Listening | Open | Closed
```

Then, we’ll begin defining an interface, starting with a `Sock`

type
for representing sockets, parameterised by their current state:

```
interface Sockets (m : Type -> Type) where
Sock : SocketState -> Type
```

We create sockets using the `socket`

method. The `SocketType`

is defined
by the sockets library, and describes whether the socket is TCP, UDP,
or some other form. We’ll use `Stream`

for this throughout, which indicates a
TCP socket.

```
socket : SocketType -> ST m (Either () Var) [addIfRight (Sock Ready)]
```

Remember that `addIfRight`

adds a resource if the result of the operation
is of the form `Right val`

. By convention in this interface, we’ll use
`Either`

for operations which might fail, whether or not they might carry
any additional information about the error, so that we can consistently
use `addIfRight`

and some other type level functions.

To define a server, once we’ve created a socket, we need to `bind`

it
to a port. We can do this with the `bind`

method:

```
bind : (sock : Var) -> (addr : Maybe SocketAddress) -> (port : Port) ->
ST m (Either () ()) [sock ::: Sock Ready :-> (Sock Closed `or` Sock Bound)]
```

Binding a socket might fail, for example if there is already a socket
bound to the given port, so again it returns a value of type `Either`

.
The action here uses a type level function `or`

, and says that:

- If
`bind`

fails, the socket moves to the`Sock Closed`

state - If
`bind`

succeeds, the socket moves to the`Sock Bound`

state, as shown in the diagram above

`or`

is implemented as follows:

```
or : a -> a -> Either b c -> a
or x y = either (const x) (const y)
```

So, the type of `bind`

could equivalently be written as:

```
bind : (sock : Var) -> (addr : Maybe SocketAddress) -> (port : Port) ->
STrans m (Either () ()) [sock ::: Sock Ready]
(either [sock ::: Sock Closed] [sock ::: Sock Bound])
```

However, using `or`

is much more concise than this, and attempts to
reflect the state transition diagram as directly as possible while still
capturing the possibility of failure.

Once we’ve bound a socket to a port, we can start listening for connections from clients:

```
listen : (sock : Var) ->
ST m (Either () ()) [sock ::: Sock Bound :-> (Sock Closed `or` Sock Listening)]
```

A socket in the `Listening`

state is ready to accept connections from
individual clients:

```
accept : (sock : Var) ->
ST m (Either () Var)
[sock ::: Sock Listening, addIfRight (Sock Open)]
```

If there is an incoming connection from a client, `accept`

adds a *new*
resource to the end of the resource list (by convention, it’s a good idea
to add resources to the end of the list, because this works more tidily
with `updateWith`

, as discussed in the previous section). So, we now
have *two* sockets: one continuing to listen for incoming connections,
and one ready for communication with the client.

We also need methods for sending and receiving data on a socket:

```
send : (sock : Var) -> String ->
ST m (Either () ()) [sock ::: Sock Open :-> (Sock Closed `or` Sock Open)]
recv : (sock : Var) ->
ST m (Either () String) [sock ::: Sock Open :-> (Sock Closed `or` Sock Open)]
```

Once we’ve finished communicating with another machine via a socket, we’ll
want to `close`

the connection and remove the socket:

```
close : (sock : Var) ->
{auto prf : CloseOK st} -> ST m () [sock ::: Sock st :-> Sock Closed]
remove : (sock : Var) ->
ST m () [Remove sock (Sock Closed)]
```

We have a predicate `CloseOK`

, used by `close`

in an implicit proof
argument, which describes when it is okay to close a socket:

```
data CloseOK : SocketState -> Type where
CloseOpen : CloseOK Open
CloseListening : CloseOK Listening
```

That is, we can close a socket which is `Open`

, talking to another machine,
which causes the communication to terminate. We can also close a socket which
is `Listening`

for incoming connections, which causes the server to stop
accepting requests.

In this section, we’re implementing a server, but for completeness we may
also want a client to connect to a server on another machine. We can do
this with `connect`

:

```
connect : (sock : Var) -> SocketAddress -> Port ->
ST m (Either () ()) [sock ::: Sock Ready :-> (Sock Closed `or` Sock Open)]
```

For reference, here is the complete interface:

```
interface Sockets (m : Type -> Type) where
Sock : SocketState -> Type
socket : SocketType -> ST m (Either () Var) [addIfRight (Sock Ready)]
bind : (sock : Var) -> (addr : Maybe SocketAddress) -> (port : Port) ->
ST m (Either () ()) [sock ::: Sock Ready :-> (Sock Closed `or` Sock Bound)]
listen : (sock : Var) ->
ST m (Either () ()) [sock ::: Sock Bound :-> (Sock Closed `or` Sock Listening)]
accept : (sock : Var) ->
ST m (Either () Var) [sock ::: Sock Listening, addIfRight (Sock Open)]
connect : (sock : Var) -> SocketAddress -> Port ->
ST m (Either () ()) [sock ::: Sock Ready :-> (Sock Closed `or` Sock Open)]
close : (sock : Var) -> {auto prf : CloseOK st} ->
ST m () [sock ::: Sock st :-> Sock Closed]
remove : (sock : Var) -> ST m () [Remove sock (Sock Closed)]
send : (sock : Var) -> String ->
ST m (Either () ()) [sock ::: Sock Open :-> (Sock Closed `or` Sock Open)]
recv : (sock : Var) ->
ST m (Either () String) [sock ::: Sock Open :-> (Sock Closed `or` Sock Open)]
```

We’ll see how to implement this shortly; mostly, the methods can be implemented
in `IO`

by using the raw sockets API directly. First, though, we’ll see
how to use the API to implement an “echo” server.

#### Implementing an “Echo” server with `Sockets`

¶

At the top level, our echo server begins and ends with no resources available,
and uses the `ConsoleIO`

and `Sockets`

interfaces:

```
startServer : (ConsoleIO m, Sockets m) => ST m () []
```

The first thing we need to do is create a socket for binding to a port
and listening for incoming connections, using `socket`

. This might fail,
so we’ll need to deal with the case where it returns `Right sock`

, where
`sock`

is the new socket variable, or where it returns `Left err`

:

```
startServer : (ConsoleIO m, Sockets m) => ST m () []
startServer =
do Right sock <- socket Stream
| Left err => pure ()
?whatNow
```

It’s a good idea to implement this kind of function interactively, step by
step, using holes to see what state the overall system is in after each
step. Here, we can see that after a successful call to `socket`

, we
have a socket available in the `Ready`

state:

```
sock : Var
m : Type -> Type
constraint : ConsoleIO m
constraint1 : Sockets m
--------------------------------------
whatNow : STrans m () [sock ::: Sock Ready] (\result1 => [])
```

Next, we need to bind the socket to a port, and start listening for
connections. Again, each of these could fail. If they do, we’ll remove
the socket. Failure always results in a socket in the `Closed`

state,
so all we can do is `remove`

it:

```
startServer : (ConsoleIO m, Sockets m) => ST m () []
startServer =
do Right sock <- socket Stream | Left err => pure ()
Right ok <- bind sock Nothing 9442 | Left err => remove sock
Right ok <- listen sock | Left err => remove sock
?runServer
```

Finally, we have a socket which is listening for incoming connections:

```
ok : ()
sock : Var
ok1 : ()
m : Type -> Type
constraint : ConsoleIO m
constraint1 : Sockets m
--------------------------------------
runServer : STrans m () [sock ::: Sock Listening]
(\result1 => [])
```

We’ll implement this in a separate function. The type of `runServer`

tells us what the type of `echoServer`

must be (noting that we need
to give the `m`

argument to `Sock`

explicitly):

```
echoServer : (ConsoleIO m, Sockets m) => (sock : Var) ->
ST m () [remove sock (Sock {m} Listening)]
```

We can complete the definition of `startServer`

as follows:

```
startServer : (ConsoleIO m, Sockets m) => ST m () []
startServer =
do Right sock <- socket Stream | Left err => pure ()
Right ok <- bind sock Nothing 9442 | Left err => remove sock
Right ok <- listen sock | Left err => remove sock
echoServer sock
```

In `echoServer`

, we’ll keep accepting requests and responding to them
until something fails, at which point we’ll close the sockets and
return. We begin by trying to accept an incoming connection:

```
echoServer : (ConsoleIO m, Sockets m) => (sock : Var) ->
ST m () [remove sock (Sock {m} Listening)]
echoServer sock =
do Right new <- accept sock | Left err => do close sock; remove sock
?whatNow
```

If `accept`

fails, we need to close the `Listening`

socket and
remove it before returning, because the type of `echoServer`

requires
this.

As always, implementing `echoServer`

incrementally means that we can check
the state we’re in as we develop. If `accept`

succeeds, we have the
existing `sock`

which is still listening for connections, and a `new`

socket, which is open for communication:

```
new : Var
sock : Var
m : Type -> Type
constraint : ConsoleIO m
constraint1 : Sockets m
--------------------------------------
whatNow : STrans m () [sock ::: Sock Listening, new ::: Sock Open]
(\result1 => [])
```

To complete `echoServer`

, we’ll receive a message on the `new`

socket, and echo it back. When we’re done, we close the `new`

socket,
and go back to the beginning of `echoServer`

to handle the next
connection:

```
echoServer : (ConsoleIO m, Sockets m) => (sock : Var) ->
ST m () [remove sock (Sock {m} Listening)]
echoServer sock =
do Right new <- accept sock | Left err => do close sock; remove sock
Right msg <- recv new | Left err => do close sock; remove sock; remove new
Right ok <- send new ("You said " ++ msg)
| Left err => do remove new; close sock; remove sock
close new; remove new; echoServer sock
```

#### Implementing `Sockets`

¶

To implement `Sockets`

in `IO`

, we’ll begin by giving a concrete type
for `Sock`

. We can use the raw sockets API (implemented in
`Network.Sockeet`

) for this, and use a `Socket`

stored in a `State`

, no
matter what abstract state the socket is in:

```
implementation Sockets IO where
Sock _ = State Socket
```

Most of the methods can be implemented by using the raw socket API
directly, returning `Left`

or `Right`

as appropriate. For example,
we can implement `socket`

, `bind`

and `listen`

as follows:

```
socket ty = do Right sock <- lift $ Socket.socket AF_INET ty 0
| Left err => pure (Left ())
lbl <- new sock
pure (Right lbl)
bind sock addr port = do ok <- lift $ bind !(read sock) addr port
if ok /= 0
then pure (Left ())
else pure (Right ())
listen sock = do ok <- lift $ listen !(read sock)
if ok /= 0
then pure (Left ())
else pure (Right ())
```

There is a small difficulty with `accept`

, however, because when we
use `new`

to create a new resource for the open connection, it appears
at the *start* of the resource list, not the end. We can see this by
writing an incomplete definition, using `returning`

to see what the
resources need to be if we return `Right lbl`

:

```
accept sock = do Right (conn, addr) <- lift $ accept !(read sock)
| Left err => pure (Left ())
lbl <- new conn
returning (Right lbl) ?fixResources
```

It’s convenient for `new`

to add the resource to the beginning of the
list because, in general, this makes automatic proof construction with
an `auto`

-implicit easier for Idris. On the other hand, when we use
`call`

to make a smaller set of resources, `updateWith`

puts newly
created resources at the *end* of the list, because in general that reduces
the amount of re-ordering of resources.

If we look at the type of
`fixResources`

, we can see what we need to do to finish `accept`

:

```
_bindApp0 : Socket
conn : Socket
addr : SocketAddress
sock : Var
lbl : Var
--------------------------------------
fixResources : STrans IO () [lbl ::: State Socket, sock ::: State Socket]
(\value => [sock ::: State Socket, lbl ::: State Socket])
```

The current list of resources is ordered `lbl`

, `sock`

, and we need them
to be in the order `sock`

, `lbl`

. To help with this situation,
`Control.ST`

provides a primitive `toEnd`

which moves a resource to the
end of the list. We can therefore complete `accept`

as follows:

```
accept sock = do Right (conn, addr) <- lift $ accept !(read sock)
| Left err => pure (Left ())
lbl <- new conn
returning (Right lbl) (toEnd lbl)
```

For the complete implementation of `Sockets`

, take a look at
`samples/ST/Net/Network.idr`

in the Idris distribution. You can also
find the complete echo server there, `EchoServer.idr`

. There is also
a higher level network protocol, `RandServer.idr`

, using a hierarchy of
state machines to implement a high level network communication protocol
in terms of the lower level sockets API. This also uses threading, to
handle incoming requests asynchronously. You can find some more detail
on threading and the random number server in the draft paper
State Machines All The Way Down
by Edwin Brady.

## The Effects Tutorial¶

A tutorial on the Effects package in Idris.

Effects and the `Control.ST`

module

There is a new module in the `contrib`

package, `Control.ST`

, which
provides the resource tracking facilities of Effects but with
better support for creating and deleting resources, and implementing
resources in terms of other resources.

Unless you have a particular reason to use Effects you are strongly
recommended to use `Control.ST`

instead. There is a tutorial available
on this site for `Control.ST`

with several examples
(Implementing State-aware Systems in Idris: The ST Tutorial).

Note

The documentation for Idris has been published under the Creative
Commons CC0 License. As such to the extent possible under law, *The
Idris Community* has waived all copyright and related or neighbouring
rights to Documentation for Idris.

### Introduction¶

Pure functional languages with dependent types such as Idris support reasoning about programs directly
in the type system, promising that we can *know* a program will run
correctly (i.e. according to the specification in its type) simply
because it compiles. Realistically, though, things are not so simple:
programs have to interact with the outside world, with user input,
input from a network, mutable state, and so on. In this tutorial I
will introduce the library, which is included with the distribution
and supports programming and reasoning with side-effecting programs,
supporting mutable state, interaction with the outside world,
exceptions, and verified resource management.

This tutorial assumes familiarity with pure programming in Idris, as described in Sections 1–6 of the main tutorial [1]. The examples presented are tested with Idris and can be found in the examples directory of the Idris repository.

Consider, for example, the following introductory function which illustrates the kind of properties which can be expressed in the type system:

```
vadd : Vect n Int -> Vect n Int -> Vect n Int
vadd [] [] = []
vadd (x :: xs) (y :: ys) = x + y :: vadd xs ys
```

This function adds corresponding elements in a pair of vectors. The type
guarantees that the vectors will contain only elements of type `Int`

,
and that the input lengths and the output length all correspond. A
natural question to ask here, which is typically neglected by
introductory tutorials, is “How do I turn this into a program?” That is,
given some lists entered by a user, how do we get into a position to be
able to apply the `vadd`

function? Before doing so, we will have to:

- Read user input, either from the keyboard, a file, or some other input device.
- Check that the user inputs are valid, i.e. contain only
`Int`

and are the same length, and report an error if not. - Write output

The complete program will include side-effects for I/O and error
handling, before we can get to the pure core functionality. In this
tutorial, we will see how Idris supports side-effects.
Furthermore, we will see how we can use the dependent type system to
*reason* about stateful and side-effecting programs. We will return to
this specific example later.

#### Hello world¶

To give an idea of how programs with effects look, here is the
ubiquitous “Hello world” program, written using the `Effects`

library:

```
module Main
import Effects
import Effect.StdIO
hello : Eff () [STDIO]
hello = putStrLn "Hello world!"
main : IO ()
main = run hello
```

As usual, the entry point is `main`

. All `main`

has to do is invoke the
`hello`

function which supports the `STDIO`

effect for console I/O, and
returns the unit value. All programs using the `Effects`

library must
`import Effects`

. The details of the `Eff`

type will be presented in the
remainder of this tutorial.

To compile and run this program, Idris needs to be told to include
the `Effects`

package, using the `-p effects`

flag (this flag is
required for all examples in this tutorial):

```
idris hello.idr -o hello -p effects
./hello Hello world!
```

#### Outline¶

The tutorial is structured as follows: first, in Section
State, we will discuss state management, describing why it
is important and introducing the `effects`

library to show how it
can be used to manage state. This section also gives an overview of
the syntax of effectful programs. Section Simple Effects then
introduces a number of other effects a program may have: I/O;
Exceptions; Random Numbers; and Non-determinism, giving examples for
each, and an extended example combining several effects in one
complete program. Section Dependent Effects introduces *dependent*
effects, showing how states and resources can be managed in
types. Section Creating New Effects shows how new effects can be
implemented. Section Example: A “Mystery Word” Guessing Game gives an extended example
showing how to implement a “mystery word” guessing game, using effects
to describe the rules of the game and ensure they are implemented
accurately. References to further reading are given in Section
Further Reading.

[1] | You do not, however, need to know what a monad is! |

### State¶

Many programs, even pure programs, can benefit from locally mutable state. For example, consider a program which tags binary tree nodes with a counter, by an inorder traversal (i.e. counting depth first, left to right). This would perform something like the following:

We can describe binary trees with the following data type `BTree`

and `testTree`

to represent the example input above:

```
data BTree a = Leaf
| Node (BTree a) a (BTree a)
testTree : BTree String
testTree = Node (Node Leaf "Jim" Leaf)
"Fred"
(Node (Node Leaf "Alice" Leaf)
"Sheila"
(Node Leaf "Bob" Leaf))
```

Then our function to implement tagging, beginning to tag with a
specific value `i`

, has the following type:

```
treeTag : (i : Int) -> BTree a -> BTree (Int, a)
```

#### First attempt¶

Naïvely, we can implement `treeTag`

by implementing a helper
function which propagates a counter, returning the result of the count
for each subtree:

```
treeTagAux : (i : Int) -> BTree a -> (Int, BTree (Int, a))
treeTagAux i Leaf = (i, Leaf)
treeTagAux i (Node l x r)
= let (i', l') = treeTagAux i l in
let x' = (i', x) in
let (i'', r') = treeTagAux (i' + 1) r in
(i'', Node l' x' r')
treeTag : (i : Int) -> BTree a -> BTree (Int, a)
treeTag i x = snd (treeTagAux i x)
```

This gives the expected result when run at the REPL prompt:

```
*TreeTag> treeTag 1 testTree
Node (Node Leaf (1, "Jim") Leaf)
(2, "Fred")
(Node (Node Leaf (3, "Alice") Leaf)
(4, "Sheila")
(Node Leaf (5, "Bob") Leaf)) : BTree (Int, String)
```

This works as required, but there are several problems when we try to
scale this to larger programs. It is error prone, because we need to
ensure that state is propagated correctly to the recursive calls (i.e.
passing the appropriate `i`

or `i’`

). It is hard to read, because
the functional details are obscured by the state propagation. Perhaps
most importantly, there is a common programming pattern here which
should be abstracted but instead has been implemented by hand. There
is local mutable state (the counter) which we have had to make
explicit.

#### Introducing `Effects`

¶

Idris provides a library, `Effects`

[3], which captures this
pattern and many others involving effectful computation [1]. An
effectful program `f`

has a type of the following form:

```
f : (x1 : a1) -> (x2 : a2) -> ... -> Eff t effs
```

That is, the return type gives the effects that `f`

supports
(`effs`

, of type `List EFFECT`

) and the type the computation
returns `t`

. So, our `treeTagAux`

helper could be written with the
following type:

```
treeTagAux : BTree a -> Eff (BTree (Int, a)) [STATE Int]
```

That is, `treeTagAux`

has access to an integer state, because the
list of available effects includes `STATE Int`

. `STATE`

is
declared as follows in the module `Effect.State`

(that is, we must
`import Effect.State`

to be able to use it):

```
STATE : Type -> EFFECT
```

It is an effect parameterised by a type (by convention, we write
effects in all capitals). The `treeTagAux`

function is an effectful
program which builds a new tree tagged with `Ints`

, and is
implemented as follows:

```
treeTagAux Leaf = pure Leaf
treeTagAux (Node l x r)
= do l' <- treeTagAux l
i <- get
put (i + 1)
r' <- treeTagAux r
pure (Node l' (i, x) r')
```

There are several remarks to be made about this implementation.
Essentially, it hides the state, which can be accessed using `get`

and updated using `put`

, but it introduces several new features.
Specifically, it uses `do`

-notation, binding variables with `<-`

,
and a `pure`

function. There is much to be said about these
features, but for our purposes, it suffices to know the following:

`do`

blocks allow effectful operations to be sequenced.`x <- e`

binds the result of an effectful operation`e`

to a- variable
`x`

. For example, in the above code,`treeTagAux l`

is an effectful operation returning`BTree (Int, a)`

, so`l’`

has type`BTree (Int, a)`

.

`pure e`

turns a pure value`e`

into the result of an effectful- operation.

The `get`

and `put`

functions read and write a state `t`

,
assuming that the `STATE t`

effect is available. They have the
following types, polymorphic in the state `t`

they manage:

```
get : Eff t [STATE t]
put : t -> Eff () [STATE t]
```

A program in `Eff`

can call any other function in `Eff`

provided
that the calling function supports at least the effects required by
the called function. In this case, it is valid for `treeTagAux`

to
call both `get`

and `put`

because all three functions support the
`STATE Int`

effect.

Programs in `Eff`

are run in some underlying *computation context*,
using the `run`

or `runPure`

function. Using `runPure`

, which
runs an effectful program in the identity context, we can write the
`treeTag`

function as follows, using `put`

to initialise the
state:

```
treeTag : (i : Int) -> BTree a -> BTree (Int, a)
treeTag i x = runPure (do put i
treeTagAux x)
```

We could also run the program in an impure context such as `IO`

,
without changing the definition of `treeTagAux`

, by using `run`

instead of `runPure`

:

```
treeTagAux : BTree a -> Eff (BTree (Int, a)) [STATE Int]
...
treeTag : (i : Int) -> BTree a -> IO (BTree (Int, a))
treeTag i x = run (do put i
treeTagAux x)
```

Note that the definition of `treeTagAux`

is exactly as before. For
reference, this complete program (including a `main`

to run it) is
shown in Listing [introprog].

```
module Main
import Effects
import Effect.State
data BTree a = Leaf
| Node (BTree a) a (BTree a)
Show a => Show (BTree a) where
show Leaf = "[]"
show (Node l x r) = "[" ++ show l ++ " "
++ show x ++ " "
++ show r ++ "]"
testTree : BTree String
testTree = Node (Node Leaf "Jim" Leaf)
"Fred"
(Node (Node Leaf "Alice" Leaf)
"Sheila"
(Node Leaf "Bob" Leaf))
treeTagAux : BTree a -> Eff (BTree (Int, a)) [STATE Int]
treeTagAux Leaf = pure Leaf
treeTagAux (Node l x r) = do l' <- treeTagAux l
i <- get
put (i + 1)
r' <- treeTagAux r
pure (Node l' (i, x) r')
treeTag : (i : Int) -> BTree a -> BTree (Int, a)
treeTag i x = runPure (do put i; treeTagAux x)
main : IO ()
main = print (treeTag 1 testTree)
```

#### Effects and Resources¶

Each effect is associated with a *resource*, which is initialised
before an effectful program can be run. For example, in the case of
`STATE Int`

the corresponding resource is the integer state itself.
The types of `runPure`

and `run`

show this (slightly simplified
here for illustrative purposes):

```
runPure : {env : Env id xs} -> Eff a xs -> a
run : Applicative m => {env : Env m xs} -> Eff a xs -> m a
```

The `env`

argument is implicit, and initialised automatically where
possible using default values given by implementations of the following
interface:

```
interface Default a where
default : a
```

Implementations of `Default`

are defined for all primitive types, and many
library types such as `List`

, `Vect`

, `Maybe`

, pairs, etc.
However, where no default value exists for a resource type (for
example, you may want a `STATE`

type for which there is no
`Default`

implementation) the resource environment can be given explicitly
using one of the following functions:

```
runPureInit : Env id xs -> Eff a xs -> a
runInit : Applicative m => Env m xs -> Eff a xs -> m a
```

To be well-typed, the environment must contain resources corresponding
exactly to the effects in `xs`

. For example, we could also have
implemented `treeTag`

by initialising the state as follows:

```
treeTag : (i : Int) -> BTree a -> BTree (Int, a)
treeTag i x = runPureInit [i] (treeTagAux x)
```

#### Labelled Effects¶

What if we have more than one state, especially more than one state of
the same type? How would `get`

and `put`

know which state they
should be referring to? For example, how could we extend the tree
tagging example such that it additionally counts the number of leaves
in the tree? One possibility would be to change the state so that it
captured both of these values, e.g.:

```
treeTagAux : BTree a -> Eff (BTree (Int, a)) [STATE (Int, Int)]
```

Doing this, however, ties the two states together throughout (as well as not indicating which integer is which). It would be nice to be able to call effectful programs which guaranteed only to access one of the states, for example. In a larger application, this becomes particularly important.

The library therefore allows effects in general to be *labelled* so
that they can be referred to explicitly by a particular name. This
allows multiple effects of the same type to be included. We can count
leaves and update the tag separately, by labelling them as follows:

```
treeTagAux : BTree a -> Eff (BTree (Int, a))
['Tag ::: STATE Int,
'Leaves ::: STATE Int]
```

The `:::`

operator allows an arbitrary label to be given to an
effect. This label can be any type—it is simply used to identify an
effect uniquely. Here, we have used a symbol type. In general
`’name`

introduces a new symbol, the only purpose of which is to
disambiguate values [2].

When an effect is labelled, its operations are also labelled using the
`:-`

operator. In this way, we can say explicitly which state we
mean when using `get`

and `put`

. The tree tagging program which
also counts leaves can be written as follows:

```
treeTagAux Leaf = do
'Leaves :- update (+1)
pure Leaf
treeTagAux (Node l x r) = do
l' <- treeTagAux l
i <- 'Tag :- get
'Tag :- put (i + 1)
r' <- treeTagAux r
pure (Node l' (i, x) r')
```

The `update`

function here is a combination of `get`

and `put`

,
applying a function to the current state.

```
update : (x -> x) -> Eff () [STATE x]
```

Finally, our top level `treeTag`

function now returns a pair of the
number of leaves, and the new tree. Resources for labelled effects are
initialised using the `:=`

operator (reminiscent of assignment in an
imperative language):

```
treeTag : (i : Int) -> BTree a -> (Int, BTree (Int, a))
treeTag i x = runPureInit ['Tag := i, 'Leaves := 0]
(do x' <- treeTagAux x
leaves <- 'Leaves :- get
pure (leaves, x'))
```

To summarise, we have:

`:::`

to convert an effect to a labelled effect.`:-`

to convert an effectful operation to a labelled effectful operation.`:=`

to initialise a resource for a labelled effect.

Or, more formally with their types (slightly simplified to account only for the situation where available effects are not updated):

```
(:::) : lbl -> EFFECT -> EFFECT
(:-) : (l : lbl) -> Eff a [x] -> Eff a [l ::: x]
(:=) : (l : lbl) -> res -> LRes l res
```

Here, `LRes`

is simply the resource type associated with a labelled
effect. Note that labels are polymorphic in the label type `lbl`

.
Hence, a label can be anything—a string, an integer, a type, etc.

`!`

-notation¶

In many cases, using `do`

-notation can make programs unnecessarily
verbose, particularly in cases where the value bound is used once,
immediately. The following program returns the length of the
`String`

stored in the state, for example:

```
stateLength : Eff Nat [STATE String]
stateLength = do x <- get
pure (length x)
```

This seems unnecessarily verbose, and it would be nice to program in a
more direct style in these cases. provides `!`

-notation to help with
this. The above program can be written instead as:

```
stateLength : Eff Nat [STATE String]
stateLength = pure (length !get)
```

The notation `!expr`

means that the expression `expr`

should be
evaluated and then implicitly bound. Conceptually, we can think of
`!`

as being a prefix function with the following type:

```
(!) : Eff a xs -> a
```

Note, however, that it is not really a function, merely syntax! In
practice, a subexpression `!expr`

will lift `expr`

as high as
possible within its current scope, bind it to a fresh name `x`

, and
replace `!expr`

with `x`

. Expressions are lifted depth first, left
to right. In practice, `!`

-notation allows us to program in a more
direct style, while still giving a notational clue as to which
expressions are effectful.

For example, the expression:

```
let y = 42 in f !(g !(print y) !x)
```

is lifted to:

```
let y = 42 in do y' <- print y
x' <- x
g' <- g y' x'
f g'
```

#### The Type `Eff`

¶

Underneath, `Eff`

is an overloaded function which translates to an
underlying type `EffM`

:

```
EffM : (m : Type -> Type) -> (t : Type)
-> (List EFFECT)
-> (t -> List EFFECT) -> Type
```

This is more general than the types we have been writing so far. It is
parameterised over an underlying computation context `m`

, a
result type `t`

as we have already seen, as well as a `List EFFECT`

and a
function type `t -> List EFFECT`

.

These additional parameters are the list of *input* effects, and a
list of *output* effects, computed from the result of an effectful
operation. That is: running an effectful program can change the set
of effects available! This is a particularly powerful idea, and we
will see its consequences in more detail later. Some examples of
operations which can change the set of available effects are:

- Updating a state containing a dependent type (for example adding an element to a vector).
- Opening a file for reading is an effect, but whether the file really
*is*open afterwards depends on whether the file was successfully opened. - Closing a file means that reading from the file should no longer be possible.

While powerful, this can make uses of the `EffM`

type hard to read.
Therefore the library provides an overloaded function `Eff`

There are the following three versions:

```
SimpleEff.Eff : (t : Type) -> (input_effs : List EFFECT) -> Type
TransEff.Eff : (t : Type) -> (input_effs : List EFFECT) ->
(output_effs : List EFFECT) -> Type
DepEff.Eff : (t : Type) -> (input_effs : List EFFECT) ->
(output_effs_fn : t -> List EFFECT) -> Type
```

So far, we have used only the first version, `SimpleEff.Eff`

, which
is defined as follows:

```
Eff : (x : Type) -> (es : List EFFECT) -> Type
Eff x es = {m : Type -> Type} -> EffM m x es (\v => es)
```

i.e. the set of effects remains the same on output. This suffices for
the `STATE`

example we have seen so far, and for many useful
side-effecting programs. We could also have written `treeTagAux`

with the expanded type:

```
treeTagAux : BTree a ->
EffM m (BTree (Int, a)) [STATE Int] (\x => [STATE Int])
```

Later, we will see programs which update effects:

```
Eff a xs xs'
```

which is expanded to

```
EffM m a xs (\_ => xs')
```

i.e. the set of effects is updated to `xs’`

(think of a transition
in a state machine). There is, for example, a version of `put`

which
updates the type of the state:

```
putM : y -> Eff () [STATE x] [STATE y]
```

Also, we have:

```
Eff t xs (\res => xs')
```

which is expanded to

```
EffM m t xs (\res => xs')
```

i.e. the set of effects is updated according to the result of the
operation `res`

, of type `t`

.

Parameterising `EffM`

over an underlying computation context allows us
to write effectful programs which are specific to one context, and in some
cases to write programs which *extend* the list of effects available using
the `new`

function, though this is beyond the scope of this tutorial.

[1] | The earlier paper [3] describes the essential implementation details, although the library presented there is an earlier version which is less powerful than that presented in this tutorial. |

[2] | In practice, `’name` simply introduces a new empty type |

[3] | (1, 2) Edwin Brady. 2013. Programming and reasoning with algebraic
effects and dependent types. SIGPLAN Not. 48, 9 (September
2013), 133-144. DOI=10.1145/2544174.2500581
http://dl.acm.org/citation.cfm?doid=2544174.2500581 |

### Simple Effects¶

So far we have seen how to write programs with locally mutable state
using the `STATE`

effect. To recap, we have the definitions below
in a module `Effect.State`

```
module Effect.State
STATE : Type -> EFFECT
get : Eff x [STATE x]
put : x -> Eff () [STATE x]
putM : y -> Eff () [STATE x] [STATE y]
update : (x -> x) -> Eff () [STATE x]
Handler State m where { ... }
```

The last line, `Handler State m where { ... }`

, means that the `STATE`

effect is usable in any computation context `m`

. That is, a program
which uses this effect and returns something of type `a`

can be
evaluated to something of type `m a`

using `run`

, for any
`m`

. The lower case `State`

is a data type describing the
operations which make up the `STATE`

effect itself—we will go into
more detail about this in Section Creating New Effects.

In this section, we will introduce some other supported effects,
allowing console I/O, exceptions, random number generation and
non-deterministic programming. For each effect we introduce, we will
begin with a summary of the effect, its supported operations, and the
contexts in which it may be used, like that above for `STATE`

, and
go on to present some simple examples. At the end, we will see some
examples of programs which combine multiple effects.

All of the effects in the library, including those described in this section, are summarised in Appendix Effects Summary.

#### Console I/O¶

Console I/O is supported with the `STDIO`

effect, which allows reading and writing characters and strings to and
from standard input and standard output. Notice that there is a
constraint here on the computation context `m`

, because it only
makes sense to support console I/O operations in a context where we
can perform (or at the very least simulate) console I/O:

```
module Effect.StdIO
STDIO : EFFECT
putChar : Char -> Eff () [STDIO]
putStr : String -> Eff () [STDIO]
putStrLn : String -> Eff () [STDIO]
getStr : Eff String [STDIO]
getChar : Eff Char [STDIO]
Handler StdIO IO where { ... }
Handler StdIO (IOExcept a) where { ... }
```

##### Examples¶

A program which reads the user’s name, then says hello, can be written as follows:

```
hello : Eff () [STDIO]
hello = do putStr "Name? "
x <- getStr
putStrLn ("Hello " ++ trim x ++ "!")
```

We use `trim`

here to remove the trailing newline from the
input. The resource associated with `STDIO`

is simply the empty
tuple, which has a default value `()`

, so we can run this as
follows:

```
main : IO ()
main = run hello
```

In `hello`

we could also use `!`

-notation instead of ```
x <-
getStr
```

, since we only use the string that is read once:

```
hello : Eff () [STDIO]
hello = do putStr "Name? "
putStrLn ("Hello " ++ trim !getStr ++ "!")
```

More interestingly, we can combine multiple effects in one program. For example, we can loop, counting the number of people we’ve said hello to:

```
hello : Eff () [STATE Int, STDIO]
hello = do putStr "Name? "
putStrLn ("Hello " ++ trim !getStr ++ "!")
update (+1)
putStrLn ("I've said hello to: " ++ show !get ++ " people")
hello
```

The list of effects given in `hello`

means that the function can
call `get`

and `put`

on an integer state, and any functions which
read and write from the console. To run this, `main`

does not need
to be changed.

##### Aside: Resource Types¶

To find out the resource type of an effect, if necessary (for example
if we want to initialise a resource explicitly with `runInit`

rather
than using a default value with `run`

) we can run the
`resourceType`

function at the REPL:

```
*ConsoleIO> resourceType STDIO
() : Type
*ConsoleIO> resourceType (STATE Int)
Int : Type
```

#### Exceptions¶

The `EXCEPTION`

effect is declared in module `Effect.Exception`

. This allows programs
to exit immediately with an error, or errors to be handled more
generally:

```
module Effect.Exception
EXCEPTION : Type -> EFFECT
raise : a -> Eff b [EXCEPTION a]
Handler (Exception a) Maybe where { ... }
Handler (Exception a) List where { ... }
Handler (Exception a) (Either a) where { ... }
Handler (Exception a) (IOExcept a) where { ... }
Show a => Handler (Exception a) IO where { ... }
```

##### Example¶

Suppose we have a `String`

which is expected to represent an integer
in the range `0`

to `n`

. We can write a function `parseNumber`

which returns an `Int`

if parsing the string returns a number in the
appropriate range, or throws an exception otherwise. Exceptions are
parameterised by an error type:

```
data Error = NotANumber | OutOfRange
parseNumber : Int -> String -> Eff Int [EXCEPTION Error]
parseNumber num str
= if all isDigit (unpack str)
then let x = cast str in
if (x >=0 && x <= num)
then pure x
else raise OutOfRange
else raise NotANumber
```

Programs which support the `EXCEPTION`

effect can be run in any
context which has some way of throwing errors, for example, we can run
`parseNumber`

in the `Either Error`

context. It returns a value of
the form `Right x`

if successful:

```
*Exception> the (Either Error Int) $ run (parseNumber 42 "20")
Right 20 : Either Error Int
```

Or `Left e`

on failure, carrying the appropriate exception:

```
*Exception> the (Either Error Int) $ run (parseNumber 42 "50")
Left OutOfRange : Either Error Int
*Exception> the (Either Error Int) $ run (parseNumber 42 "twenty")
Left NotANumber : Either Error Int
```

In fact, we can do a little bit better with `parseNumber`

, and have
it return a *proof* that the integer is in the required range along
with the integer itself. One way to do this is define a type of
bounded integers, `Bounded`

:

```
Bounded : Int -> Type
Bounded x = (n : Int ** So (n >= 0 && n <= x))
```

Recall that `So`

is parameterised by a `Bool`

, and only ```
So
True
```

is inhabited. We can use `choose`

to construct such a value
from the result of a dynamic check:

```
data So : Bool -> Type = Oh : So True
choose : (b : Bool) -> Either (So b) (So (not b))
```

We then write `parseNumber`

using `choose`

rather than an
`if/then/else`

construct, passing the proof it returns on success as
the boundedness proof:

```
parseNumber : (x : Int) -> String -> Eff (Bounded x) [EXCEPTION Error]
parseNumber x str
= if all isDigit (unpack str)
then let num = cast str in
case choose (num >=0 && num <= x) of
Left p => pure (num ** p)
Right p => raise OutOfRange
else raise NotANumber
```

#### Random Numbers¶

Random number generation is also implemented by the library, in module
`Effect.Random`

:

```
module Effect.Random
RND : EFFECT
srand : Integer -> Eff () [RND]
rndInt : Integer -> Integer -> Eff Integer [RND]
rndFin : (k : Nat) -> Eff (Fin (S k)) [RND]
Handler Random m where { ... }
```

Random number generation is considered side-effecting because its
implementation generally relies on some external source of randomness.
The default implementation here relies on an integer *seed*, which can
be set with `srand`

. A specific seed will lead to a predictable,
repeatable sequence of random numbers. There are two functions which
produce a random number:

`rndInt`

, which returns a random integer between the given lower- and upper bounds.

`rndFin`

, which returns a random element of a finite set- (essentially a number with an upper bound given in its type).

##### Example¶

We can use the `RND`

effect to implement a simple guessing game. The
`guess`

function, given a target number, will repeatedly ask the
user for a guess, and state whether the guess is too high, too low, or
correct:

```
guess : Int -> Eff () [STDIO]
```

For reference, the code for `guess`

is given below:

```
guess : Int -> Eff () [STDIO]
guess target
= do putStr "Guess: "
case run {m=Maybe} (parseNumber 100 (trim !getStr)) of
Nothing => do putStrLn "Invalid input"
guess target
Just (v ** _) =>
case compare v target of
LT => do putStrLn "Too low"
guess target
EQ => putStrLn "You win!"
GT => do putStrLn "Too high"
guess target
```

Note that we use `parseNumber`

as defined previously to read user input, but
we don’t need to list the `EXCEPTION`

effect because we use a nested `run`

to invoke `parseNumber`

, independently of the calling effectful program.

To invoke this, we pick a random number within the range 0–100,
having set up the random number generator with a seed, then run
`guess`

:

```
game : Eff () [RND, STDIO]
game = do srand 123456789
guess (fromInteger !(rndInt 0 100))
main : IO ()
main = run game
```

If no seed is given, it is set to the `default`

value. For a less
predictable game, some better source of randomness would be required,
for example taking an initial seed from the system time. To see how to
do this, see the `SYSTEM`

effect described in Effects Summary.

#### Non-determinism¶

The listing below gives the definition of the non-determinism effect, which allows a program to choose a value non-deterministically from a list of possibilities in such a way that the entire computation succeeds:

```
import Effects
import Effect.Select
SELECT : EFFECT
select : List a -> Eff a [SELECT]
Handler Selection Maybe where { ... }
Handler Selection List where { ... }
```

##### Example¶

The `SELECT`

effect can be used to solve constraint problems, such
as finding Pythagorean triples. The idea is to use `select`

to give
a set of candidate values, then throw an exception for any combination
of values which does not satisfy the constraint:

```
triple : Int -> Eff (Int, Int, Int) [SELECT, EXCEPTION String]
triple max = do z <- select [1..max]
y <- select [1..z]
x <- select [1..y]
if (x * x + y * y == z * z)
then pure (x, y, z)
else raise "No triple"
```

This program chooses a value for `z`

between `1`

and `max`

, then
values for `y`

and `x`

. In operation, after a `select`

, the
program executes the rest of the `do`

-block for every possible
assignment, effectively searching depth-first. If the list is empty
(or an exception is thrown) execution fails.

There are handlers defined for `Maybe`

and `List`

contexts, i.e.
contexts which can capture failure. Depending on the context `m`

,
`triple`

will either return the first triple it finds (if in
`Maybe`

context) or all triples in the range (if in `List`

context). We can try this as follows:

```
main : IO ()
main = do print $ the (Maybe _) $ run (triple 100)
print $ the (List _) $ run (triple 100)
```

`vadd`

revisited¶

We now return to the `vadd`

program from the introduction. Recall the
definition:

```
vadd : Vect n Int -> Vect n Int -> Vect n Int
vadd [] [] = []
vadd (x :: xs) (y :: ys) = x + y :: vadd xs ys
```

Using , we can set up a program so that it reads input from a user,
checks that the input is valid (i.e both vectors contain integers, and
are the same length) and if so, pass it on to `vadd`

. First, we
write a wrapper for `vadd`

which checks the lengths and throw an
exception if they are not equal. We can do this for input vectors of
length `n`

and `m`

by matching on the implicit arguments `n`

and
`m`

and using `decEq`

to produce a proof of their equality, if
they are equal:

```
vadd_check : Vect n Int -> Vect m Int ->
Eff (Vect m Int) [EXCEPTION String]
vadd_check {n} {m} xs ys with (decEq n m)
vadd_check {n} {m=n} xs ys | (Yes Refl) = pure (vadd xs ys)
vadd_check {n} {m} xs ys | (No contra) = raise "Length mismatch"
```

To read a vector from the console, we implement a function of the following type:

```
read_vec : Eff (p ** Vect p Int) [STDIO]
```

This returns a dependent pair of a length, and a vector of that
length, because we cannot know in advance how many integers the user
is going to input. We can use `-1`

to indicate the end of input:

```
read_vec : Eff (p ** Vect p Int) [STDIO]
read_vec = do putStr "Number (-1 when done): "
case run (parseNumber (trim !getStr)) of
Nothing => do putStrLn "Input error"
read_vec
Just v => if (v /= -1)
then do (_ ** xs) <- read_vec
pure (_ ** v :: xs)
else pure (_ ** [])
where
parseNumber : String -> Eff Int [EXCEPTION String]
parseNumber str
= if all (\x => isDigit x || x == '-') (unpack str)
then pure (cast str)
else raise "Not a number"
```

This uses a variation on `parseNumber`

which does not require a
number to be within range.

Finally, we write a program which reads two vectors and prints the result of pairwise addition of them, throwing an exception if the inputs are of differing lengths:

```
do_vadd : Eff () [STDIO, EXCEPTION String]
do_vadd = do putStrLn "Vector 1"
(_ ** xs) <- read_vec
putStrLn "Vector 2"
(_ ** ys) <- read_vec
putStrLn (show !(vadd_check xs ys))
```

By having explicit lengths in the type, we can be sure that `vadd`

is only being used where the lengths of inputs are guaranteed to be
equal. This does not stop us reading vectors from user input, but it
does require that the lengths are checked and any discrepancy is dealt
with gracefully.

#### Example: An Expression Calculator¶

To show how these effects can fit together, let us consider an evaluator for a simple expression language, with addition and integer values.

```
data Expr = Val Integer
| Add Expr Expr
```

An evaluator for this language always returns an `Integer`

, and
there are no situations in which it can fail!

```
eval : Expr -> Integer
eval (Val x) = x
eval (Add l r) = eval l + eval r
```

If we add variables, however, things get more interesting. The evaluator will need to be able to access the values stored in variables, and variables may be undefined.

```
data Expr = Val Integer
| Var String
| Add Expr Expr
```

To start, we will change the type of `eval`

so that it is effectful,
and supports an exception effect for throwing errors, and a state
containing a mapping from variable names (as `String`

) to their
values:

```
Env : Type
Env = List (String, Integer)
eval : Expr -> Eff Integer [EXCEPTION String, STATE Env]
eval (Val x) = pure x
eval (Add l r) = pure $ !(eval l) + !(eval r)
```

Note that we are using `!`

-notation to avoid having to bind
subexpressions in a `do`

block. Next, we add a case for evaluating
`Var`

:

```
eval (Var x) = case lookup x !get of
Nothing => raise $ "No such variable " ++ x
Just val => pure val
```

This retrieves the state (with `get`

, supported by the `STATE Env`

effect) and raises an exception if the variable is not in the
environment (with `raise`

, supported by the `EXCEPTION String`

effect).

To run the evaluator on a particular expression in a particular
environment of names and their values, we can write a function which
sets the state then invokes `eval`

:

```
runEval : List (String, Integer) -> Expr -> Maybe Integer
runEval args expr = run (eval' expr)
where eval' : Expr -> Eff Integer [EXCEPTION String, STATE Env]
eval' e = do put args
eval e
```

We have picked `Maybe`

as a computation context here; it needs to be
a context which is available for every effect supported by
`eval`

. In particular, because we have exceptions, it needs to be a
context which supports exceptions. Alternatively, `Either String`

or
`IO`

would be fine, for example.

What if we want to extend the evaluator further, with random number
generation? To achieve this, we add a new constructor to `Expr`

,
which gives a random number up to a maximum value:

```
data Expr = Val Integer
| Var String
| Add Expr Expr
| Random Integer
```

Then, we need to deal with the new case, making sure that we extend
the list of events to include `RND`

. It doesn’t matter where `RND`

appears in the list, as long as it is present:

```
eval : Expr -> Eff Integer [EXCEPTION String, RND, STATE Env]
eval (Random upper) = rndInt 0 upper
```

For test purposes, we might also want to print the random number which has been generated:

```
eval (Random upper) = do val <- rndInt 0 upper
putStrLn (show val)
pure val
```

If we try this without extending the effects list, we would see an error something like the following:

```
Expr.idr:28:6:When elaborating right hand side of eval:
Can't solve goal
SubList [STDIO]
[(EXCEPTION String), RND, (STATE (List (String, Integer)))]
```

In other words, the `STDIO`

effect is not available. We can correct
this simply by updating the type of `eval`

to include `STDIO`

.

```
eval : Expr -> Eff Integer [STDIO, EXCEPTION String, RND, STATE Env]
```

Note

Using `STDIO`

will restrict the number of contexts in
which `eval`

can be `run`

to those which support
`STDIO`

, such as `IO`

. Once effect lists get longer, it
can be a good idea instead to encapsulate sets of effects in
a type synonym. This is achieved as follows, simply by
defining a function which computes a type, since types are
first class in Idris:

```
EvalEff : Type -> Type
EvalEff t = Eff t [STDIO, EXCEPTION String, RND, STATE Env]
eval : Expr -> EvalEff Integer
```

### Dependent Effects¶

In the programs we have seen so far, the available effects have remained
constant. Sometimes, however, an operation can *change* the available
effects. The simplest example occurs when we have a state with a
dependent type—adding an element to a vector also changes its type, for
example, since its length is explicit in the type. In this section, we
will see how the library supports this. Firstly, we will see how states
with dependent types can be implemented. Secondly, we will see how the
effects can depend on the *result* of an effectful operation. Finally,
we will see how this can be used to implement a type-safe and
resource-safe protocol for file management.

#### Dependent States¶

Suppose we have a function which reads input from the console, converts
it to an integer, and adds it to a list which is stored in a `STATE`

.
It might look something like the following:

```
readInt : Eff () [STATE (List Int), STDIO]
readInt = do let x = trim !getStr
put (cast x :: !get)
```

But what if, instead of a list of integers, we would like to store a
`Vect`

, maintaining the length in the type?

```
readInt : Eff () [STATE (Vect n Int), STDIO]
readInt = do let x = trim !getStr
put (cast x :: !get)
```

This will not type check! Although the vector has length `n`

on entry
to `readInt`

, it has length `S n`

on exit. The library allows us to
express this as follows:

```
readInt : Eff ()[STATE (Vect n Int), STDIO]
[STATE (Vect (S n) Int), STDIO]
readInt = do let x = trim !getStr
putM (cast x :: !get)
```

The type `Eff a xs xs'`

means that the operation
begins with effects `xs`

available, and ends with effects `xs’`

available. We have used `putM`

to update the state, where the `M`

suffix indicates that the *type* is being updated as well as the value.
It has the following type:

```
putM : y -> Eff () [STATE x] [STATE y]
```

#### Result-dependent Effects¶

Often, whether a state is updated could depend on the success or
otherwise of an operation. In our `readInt`

example, we might wish to
update the vector only if the input is a valid integer (i.e. all
digits). As a first attempt, we could try the following, returning a
`Bool`

which indicates success:

```
readInt : Eff Bool [STATE (Vect n Int), STDIO]
[STATE (Vect (S n) Int), STDIO]
readInt = do let x = trim !getStr
case all isDigit (unpack x) of
False => pure False
True => do putM (cast x :: !get)
pure True
```

Unfortunately, this will not type check because the vector does not get
extended in both branches of the `case`

!

```
MutState.idr:18:19:When elaborating right hand side of Main.case
block in readInt:
Unifying n and S n would lead to infinite value
```

Clearly, the size of the resulting vector depends on whether or not the value read from the user was valid. We can express this in the type:

```
readInt : Eff Bool [STATE (Vect n Int), STDIO]
(\ok => if ok then [STATE (Vect (S n) Int), STDIO]
else [STATE (Vect n Int), STDIO])
readInt = do let x = trim !getStr
case all isDigit (unpack x) of
False => pureM False
True => do putM (cast x :: !get)
pureM True
```

Using `pureM`

rather than `pure`

allows the output effects to be
calculated from the value given. Its type is:

```
pureM : (val : a) -> EffM m a (xs val) xs
```

When using `readInt`

, we will have to check its return
value in order to know what the new set of effects is. For example, to
read a set number of values into a vector, we could write the following:

```
readN : (n : Nat) ->
Eff () [STATE (Vect m Int), STDIO]
[STATE (Vect (n + m) Int), STDIO]
readN Z = pure ()
readN {m} (S k) = case !readInt of
True => rewrite plusSuccRightSucc k m in readN k
False => readN (S k)
```

The `case`

analysis on the result of `readInt`

means that we know in
each branch whether reading the integer succeeded, and therefore how
many values still need to be read into the vector. What this means in
practice is that the type system has verified that a necessary dynamic
check (i.e. whether reading a value succeeded) has indeed been done.

Note

Only `case`

will work here. We cannot use `if/then/else`

because the `then`

and `else`

branches must have the same
type. The `case`

construct, however, abstracts over the value
being inspected in the type of each branch.

#### File Management¶

A practical use for dependent effects is in specifying resource usage protocols and verifying that they are executed correctly. For example, file management follows a resource usage protocol with the following (informally specified) requirements:

- It is necessary to open a file for reading before reading it
- Opening may fail, so the programmer should check whether opening was successful
- A file which is open for reading must not be written to, and vice versa
- When finished, an open file handle should be closed
- When a file is closed, its handle should no longer be used

These requirements can be expressed formally in , by creating a
`FILE_IO`

effect parameterised over a file handle state, which is
either empty, open for reading, or open for writing. The `FILE_IO`

effect’s definition is given below. Note that this
effect is mainly for illustrative purposes—typically we would also like
to support random access files and better reporting of error conditions.

```
module Effect.File
import Effects
import Control.IOExcept
FILE_IO : Type -> EFFECT
data OpenFile : Mode -> Type
open : (fname : String)
-> (m : Mode)
-> Eff Bool [FILE_IO ()]
(\res => [FILE_IO (case res of
True => OpenFile m
False => ())])
close : Eff () [FILE_IO (OpenFile m)] [FILE_IO ()]
readLine : Eff String [FILE_IO (OpenFile Read)]
writeLine : String -> Eff () [FILE_IO (OpenFile Write)]
eof : Eff Bool [FILE_IO (OpenFile Read)]
Handler FileIO IO where { ... }
```

In particular, consider the type of `open`

:

```
open : (fname : String)
-> (m : Mode)
-> Eff Bool [FILE_IO ()]
(\res => [FILE_IO (case res of
True => OpenFile m
False => ())])
```

This returns a `Bool`

which indicates whether opening the file was
successful. The resulting state depends on whether the operation was
successful; if so, we have a file handle open for the stated purpose,
and if not, we have no file handle. By `case`

analysis on the result,
we continue the protocol accordingly.

```
readFile : Eff (List String) [FILE_IO (OpenFile Read)]
readFile = readAcc [] where
readAcc : List String -> Eff (List String) [FILE_IO (OpenFile Read)]
readAcc acc = if (not !eof)
then readAcc (!readLine :: acc)
else pure (reverse acc)
```

Given a function `readFile`

, above, which reads from
an open file until reaching the end, we can write a program which opens
a file, reads it, then displays the contents and closes it, as follows,
correctly following the protocol:

```
dumpFile : String -> Eff () [FILE_IO (), STDIO]
dumpFile name = case !(open name Read) of
True => do putStrLn (show !readFile)
close
False => putStrLn ("Error!")
```

The type of `dumpFile`

, with `FILE_IO ()`

in its effect list,
indicates that any use of the file resource will follow the protocol
correctly (i.e. it both begins and ends with an empty resource). If we
fail to follow the protocol correctly (perhaps by forgetting to close
the file, failing to check that `open`

succeeded, or opening the file
for writing) then we will get a compile-time error. For example,
changing `open name Read`

to `open name Write`

yields a compile-time
error of the following form:

```
FileTest.idr:16:18:When elaborating right hand side of Main.case
block in testFile:
Can't solve goal
SubList [(FILE_IO (OpenFile Read))]
[(FILE_IO (OpenFile Write)), STDIO]
```

In other words: when reading a file, we need a file which is open for
reading, but the effect list contains a `FILE_IO`

effect carrying a
file open for writing.

#### Pattern-matching bind¶

It might seem that having to test each potentially failing operation
with a `case`

clause could lead to ugly code, with lots of
nested case blocks. Many languages support exceptions to improve this,
but unfortunately exceptions may not allow completely clean resource
management—for example, guaranteeing that any `open`

which did succeed
has a corresponding close.

Idris supports *pattern-matching* bindings, such as the following:

```
dumpFile : String -> Eff () [FILE_IO (), STDIO]
dumpFile name = do True <- open name Read
putStrLn (show !readFile)
close
```

This also has a problem: we are no longer dealing with the case where opening a file failed! The solution is to extend the pattern-matching binding syntax to give brief clauses for failing matches. Here, for example, we could write:

```
dumpFile : String -> Eff () [FILE_IO (), STDIO]
dumpFile name = do True <- open name Read | False => putStrLn "Error"
putStrLn (show !readFile)
close
```

This is exactly equivalent to the definition with the explicit `case`

.
In general, in a `do`

-block, the syntax:

```
do pat <- val | <alternatives>
p
```

is desugared to

```
do x <- val
case x of
pat => p
<alternatives>
```

There can be several `alternatives`

, separated by a vertical bar
`|`

. For example, there is a `SYSTEM`

effect which supports
reading command line arguments, among other things (see Appendix
Effects Summary). To read command line arguments, we can use
`getArgs`

:

```
getArgs : Eff (List String) [SYSTEM]
```

A main program can read command line arguments as follows, where in the
list which is returned, the first element `prog`

is the executable
name and the second is an expected argument:

```
emain : Eff () [SYSTEM, STDIO]
emain = do [prog, arg] <- getArgs
putStrLn $ "Argument is " ++ arg
{- ... rest of function ... -}
```

Unfortunately, this will not fail gracefully if no argument is given, or if too many arguments are given. We can use pattern matching bind alternatives to give a better (more informative) error:

```
emain : Eff () [SYSTEM, STDIO]
emain = do [prog, arg] <- getArgs | [] => putStrLn "Can't happen!"
| [prog] => putStrLn "No arguments!"
| _ => putStrLn "Too many arguments!"
putStrLn $ "Argument is " ++ arg
{- ... rest of function ... -}
```

If `getArgs`

does not return something of the form `[prog, arg]`

the
alternative which does match is executed instead, and that value
returned.

### Creating New Effects¶

We have now seen several side-effecting operations provided by the
`Effects`

library, and examples of their use in Section
Simple Effects. We have also seen how operations may *modify*
the available effects by changing state in Section
Dependent Effects. We have not, however, yet seen how these
operations are implemented. In this section, we describe how a
selection of the available effects are implemented, and show how new
effectful operations may be provided.

#### State¶

Effects are described by *algebraic data types*, where the
constructors describe the operations provided when the effect is
available. Stateful operations are described as follows:

```
data State : Effect where
Get : State a a (\x => a)
Put : b -> State () a (\x => b)
```

`Effect`

itself is a type synonym, giving the required type for an
effect signature:

```
Effect : Type
Effect = (result : Type) ->
(input_resource : Type) ->
(output_resource : result -> Type) -> Type
```

Each effect is associated with a *resource*. The second argument to
an effect signature is the resource type on *input* to an operation,
and the third is a function which computes the resource type on
*output*. Here, it means:

`Get`

takes no arguments. It has a resource of type`a`

, which is not updated, and running the`Get`

operation returns something of type`a`

.`Put`

takes a`b`

as an argument. It has a resource of type`a`

on input, which is updated to a resource of type`b`

. Running the`Put`

operation returns the element of the unit type.

The effects library provides an overloaded function `sig`

which can make effect signatures more concise, particularly when the
result has no effect on the resource type. For `State`

, we can
write:

```
data State : Effect where
Get : sig State a a
Put : b -> sig State () a b
```

There are four versions of `sig`

, depending on whether we
are interested in the resource type, and whether we are updating the
resource. Idris will infer the appropriate version from usage.

```
NoResourceEffect.sig : Effect -> Type -> Type
NoUpdateEffect.sig : Effect -> (ret : Type) ->
(resource : Type) -> Type
UpdateEffect.sig : Effect -> (ret : Type) ->
(resource_in : Type) ->
(resource_out : Type) -> Type
DepUpdateEffect.sig : Effect -> (ret : Type) ->
(resource_in : Type) ->
(resource_out : ret -> Type) -> Type
```

In order to convert `State`

(of type `Effect`

) into something
usable in an effects list, of type `EFFECT`

, we write the following:

```
STATE : Type -> EFFECT
STATE t = MkEff t State
```

`MkEff`

constructs an `EFFECT`

by taking the resource type (here,
the `t`

which parameterises `STATE`

) and the effect signature
(here, `State`

). For reference, `EFFECT`

is declared as follows:

```
data EFFECT : Type where
MkEff : Type -> Effect -> EFFECT
```

Recall that to run an effectful program in `Eff`

, we use one of the
`run`

family of functions to run the program in a particular
computation context `m`

. For each effect, therefore, we must explain
how it is executed in a particular computation context for `run`

to
work in that context. This is achieved with the following interface:

```
interface Handler (e : Effect) (m : Type -> Type) where
handle : resource -> (eff : e t resource resource') ->
((x : t) -> resource' x -> m a) -> m a
```

We have already seen some implementation declarations in the effect
summaries in Section Simple Effects. An implementation of ```
Handler e
m
```

means that the effect declared with signature `e`

can be run in
computation context `m`

. The `handle`

function takes:

- The
`resource`

on input (so, the current value of the state for`State`

) - The effectful operation (either
`Get`

or`Put x`

for`State`

) - A
*continuation*, which we conventionally call`k`

, and should be passed the result value of the operation, and an updated resource.

There are two reasons for taking a continuation here: firstly, this is neater because there are multiple return values (a new resource and the result of the operation); secondly, and more importantly, the continuation can be called zero or more times.

A `Handler`

for `State`

simply passes on the value of the state,
in the case of `Get`

, or passes on a new state, in the case of
`Put`

. It is defined the same way for all computation contexts:

```
Handler State m where
handle st Get k = k st st
handle st (Put n) k = k () n
```

This gives enough information for `Get`

and `Put`

to be used
directly in `Eff`

programs. It is tidy, however, to define top level
functions in `Eff`

, as follows:

```
get : Eff x [STATE x]
get = call Get
put : x -> Eff () [STATE x]
put val = call (Put val)
putM : y -> Eff () [STATE x] [STATE y]
putM val = call (Put val)
```

**An implementation detail (aside):** The `call`

function converts
an `Effect`

to a function in `Eff`

, given a proof that the effect
is available. This proof can be constructed automatically by , since
it is essentially an index into a statically known list of effects:

```
call : {e : Effect} ->
(eff : e t a b) -> {auto prf : EffElem e a xs} ->
Eff t xs (\v => updateResTy v xs prf eff)
```

This is the reason for the `Can’t solve goal`

error when an effect
is not available: the implicit proof `prf`

has not been solved
automatically because the required effect is not in the list of
effects `xs`

.

Such details are not important for using the library, or even writing new effects, however.

##### Summary¶

The following listing summarises what is required to define the
`STATE`

effect:

```
data State : Effect where
Get : sig State a a
Put : b -> sig State () a b
STATE : Type -> EFFECT
STATE t = MkEff t State
Handler State m where
handle st Get k = k st st
handle st (Put n) k = k () n
get : Eff x [STATE x]
get = call Get
put : x -> Eff () [STATE x]
put val = call (Put val)
putM : y -> Eff () [STATE x] [STATE y]
putM val = call (Put val)
```

#### Console I/O¶

Then listing below gives the definition of the `STDIO`

effect, including handlers for `IO`

and `IOExcept`

. We omit the
definition of the top level `Eff`

functions, as this merely invoke
the effects `PutStr`

, `GetStr`

, `PutCh`

and `GetCh`

directly.

Note that in this case, the resource is the unit type in every case,
since the handlers merely apply the `IO`

equivalents of the effects
directly.

```
data StdIO : Effect where
PutStr : String -> sig StdIO ()
GetStr : sig StdIO String
PutCh : Char -> sig StdIO ()
GetCh : sig StdIO Char
Handler StdIO IO where
handle () (PutStr s) k = do putStr s; k () ()
handle () GetStr k = do x <- getLine; k x ()
handle () (PutCh c) k = do putChar c; k () ()
handle () GetCh k = do x <- getChar; k x ()
Handler StdIO (IOExcept a) where
handle () (PutStr s) k = do ioe_lift $ putStr s; k () ()
handle () GetStr k = do x <- ioe_lift $ getLine; k x ()
handle () (PutCh c) k = do ioe_lift $ putChar c; k () ()
handle () GetCh k = do x <- ioe_lift $ getChar; k x ()
STDIO : EFFECT
STDIO = MkEff () StdIO
```

#### Exceptions¶

The listing below gives the definition of the `Exception`

effect, including two of its handlers for `Maybe`

and `List`

. The
only operation provided is `Raise`

. The key point to note in the
definitions of these handlers is that the continuation `k`

is not
used. Running `Raise`

therefore means that computation stops with an
error.

```
data Exception : Type -> Effect where
Raise : a -> sig (Exception a) b
Handler (Exception a) Maybe where
handle _ (Raise e) k = Nothing
Handler (Exception a) List where
handle _ (Raise e) k = []
EXCEPTION : Type -> EFFECT
EXCEPTION t = MkEff () (Exception t)
```

#### Non-determinism¶

The following listing gives the definition of the `Select`

effect for writing non-deterministic programs, including a handler for
`List`

context which returns all possible successful values, and a
handler for `Maybe`

context which returns the first successful
value.

```
data Selection : Effect where
Select : List a -> sig Selection a
Handler Selection Maybe where
handle _ (Select xs) k = tryAll xs where
tryAll [] = Nothing
tryAll (x :: xs) = case k x () of
Nothing => tryAll xs
Just v => Just v
Handler Selection List where
handle r (Select xs) k = concatMap (\x => k x r) xs
SELECT : EFFECT
SELECT = MkEff () Selection
```

Here, the continuation is called multiple times in each handler, for
each value in the list of possible values. In the `List`

handler, we
accumulate all successful results, and in the `Maybe`

handler we try
the first value in the list, and try later values only if that fails.

#### File Management¶

Result-dependent effects are no different from non-dependent effects
in the way they are implemented. The listing below
illustrates this for the `FILE_IO`

effect. The syntax for state
transitions `{ x ==> {res} x’ }`

, where the result state `x’`

is
computed from the result of the operation `res`

, follows that for
the equivalent `Eff`

programs.

```
data FileIO : Effect where
Open : (fname: String)
-> (m : Mode)
-> sig FileIO Bool () (\res => case res of
True => OpenFile m
False => ())
Close : sig FileIO () (OpenFile m)
ReadLine : sig FileIO String (OpenFile Read)
WriteLine : String -> sig FileIO () (OpenFile Write)
EOF : sig FileIO Bool (OpenFile Read)
Handler FileIO IO where
handle () (Open fname m) k = do h <- openFile fname m
if !(validFile h)
then k True (FH h)
else k False ()
handle (FH h) Close k = do closeFile h
k () ()
handle (FH h) ReadLine k = do str <- fread h
k str (FH h)
handle (FH h) (WriteLine str) k = do fwrite h str
k () (FH h)
handle (FH h) EOF k = do e <- feof h
k e (FH h)
FILE_IO : Type -> EFFECT
FILE_IO t = MkEff t FileIO
```

Note that in the handler for `Open`

, the types passed to the
continuation `k`

are different depending on whether the result is
`True`

(opening succeeded) or `False`

(opening failed). This uses
`validFile`

, defined in the `Prelude`

, to test whether a file
handler refers to an open file or not.

### Example: A “Mystery Word” Guessing Game¶

In this section, we will use the techniques and specific effects discussed in the tutorial so far to implement a larger example, a simple text-based word-guessing game. In the game, the computer chooses a word, which the player must guess letter by letter. After a limited number of wrong guesses, the player loses [1].

We will implement the game by following these steps:

- Define the game state, in enough detail to express the rules
- Define the rules of the game (i.e. what actions the player may take, and how these actions affect the game state)
- Implement the rules of the game (i.e. implement state updates for each action)
- Implement a user interface which allows a player to direct actions

Step 2 may be achieved by defining an effect which depends on the state
defined in step 1. Then step 3 involves implementing a `Handler`

for
this effect. Finally, step 4 involves implementing a program in `Eff`

using the newly defined effect (and any others required to implement the
interface).

#### Step 1: Game State¶

First, we categorise the game states as running games (where there are a number of guesses available, and a number of letters still to guess), or non-running games (i.e. games which have not been started, or games which have been won or lost).

```
data GState = Running Nat Nat | NotRunning
```

Notice that at this stage, we say nothing about what it means to make a guess, what the word to be guessed is, how to guess letters, or any other implementation detail. We are only interested in what is necessary to describe the game rules.

We will, however, parameterise a concrete game state `Mystery`

over
this data:

```
data Mystery : GState -> Type
```

#### Step 2: Game Rules¶

We describe the game rules as a dependent effect, where each action has
a *precondition* (i.e. what the game state must be before carrying out
the action) and a *postcondition* (i.e. how the action affects the game
state). Informally, these actions with the pre- and postconditions are:

- Guess
Guess a letter in the word.

- Precondition: The game must be running, and there must be both guesses still available, and letters still to be guessed.
- Postcondition: If the guessed letter is in the word and not yet guessed, reduce the number of letters, otherwise reduce the number of guesses.

- Won
Declare victory

- Precondition: The game must be running, and there must be no letters still to be guessed.
- Postcondition: The game is no longer running.

- Lost
Accept defeat

- Precondition: The game must be running, and there must be no guesses left.
- Postcondition: The game is no longer running.

- NewWord
Set a new word to be guessed

- Precondition: The game must not be running.
- Postcondition: The game is running, with 6 guesses available (the choice of 6 is somewhat arbitrary here) and the number of unique letters in the word still to be guessed.

- Get
- Get a string representation of the game state. This is for display purposes; there are no pre- or postconditions.

We can make these rules precise by declaring them more formally in an effect signature:

```
data MysteryRules : Effect where
Guess : (x : Char) ->
sig MysteryRules Bool
(Mystery (Running (S g) (S w)))
(\inword => if inword
then Mystery (Running (S g) w)
else Mystery (Running g (S w)))
Won : sig MysteryRules () (Mystery (Running g 0))
(Mystery NotRunning)
Lost : sig MysteryRules () (Mystery (Running 0 g))
(Mystery NotRunning)
NewWord : (w : String) ->
sig MysteryRules () (Mystery NotRunning) (Mystery (Running 6 (length (letters w))))
Get : sig MysteryRules String (Mystery h)
```

This description says nothing about how the rules are implemented. In
particular, it does not specify *how* to tell whether a guessed letter
was in a word, just that the result of `Guess`

depends on it.

Nevertheless, we can still create an `EFFECT`

from this, and use it in
an `Eff`

program. Implementing a `Handler`

for `MysteryRules`

will
then allow us to play the game.

```
MYSTERY : GState -> EFFECT
MYSTERY h = MkEff (Mystery h) MysteryRules
```

#### Step 3: Implement Rules¶

To *implement* the rules, we begin by giving a concrete definition of
game state:

```
data Mystery : GState -> Type where
Init : Mystery NotRunning
GameWon : (word : String) -> Mystery NotRunning
GameLost : (word : String) -> Mystery NotRunning
MkG : (word : String) ->
(guesses : Nat) ->
(got : List Char) ->
(missing : Vect m Char) ->
Mystery (Running guesses m)
```

If a game is `NotRunning`

, that is either because it has not yet
started (`Init`

) or because it is won or lost (`GameWon`

and
`GameLost`

, each of which carry the word so that showing the game
state will reveal the word to the player). Finally, `MkG`

captures a
running game’s state, including the target word, the letters
successfully guessed, and the missing letters. Using a `Vect`

for the
missing letters is convenient since its length is used in the type.

To initialise the state, we implement the following functions:
`letters`

, which returns a list of unique letters in a `String`

(ignoring spaces) and `initState`

which sets up an initial state
considered valid as a postcondition for `NewWord`

.

```
letters : String -> List Char
initState : (x : String) -> Mystery (Running 6 (length (letters x)))
```

When checking if a guess is in the vector of missing letters, it is
convenient to return a *proof* that the guess is in the vector, using
`isElem`

below, rather than merely a `Bool`

:

```
data IsElem : a -> Vect n a -> Type where
First : IsElem x (x :: xs)
Later : IsElem x xs -> IsElem x (y :: xs)
isElem : DecEq a => (x : a) -> (xs : Vect n a) -> Maybe (IsElem x xs)
```

The reason for returning a proof is that we can use it to remove an element from the correct position in a vector:

```
shrink : (xs : Vect (S n) a) -> IsElem x xs -> Vect n a
```

We leave the definitions of `letters`

, `init`

, `isElem`

and
`shrink`

as exercises. Having implemented these, the `Handler`

implementation for `MysteryRules`

is surprisingly straightforward:

```
Handler MysteryRules m where
handle (MkG w g got []) Won k = k () (GameWon w)
handle (MkG w Z got m) Lost k = k () (GameLost w)
handle st Get k = k (show st) st
handle st (NewWord w) k = k () (initState w)
handle (MkG w (S g) got m) (Guess x) k =
case isElem x m of
Nothing => k False (MkG w _ got m)
(Just p) => k True (MkG w _ (x :: got) (shrink m p))
```

Each case simply involves directly updating the game state in a way
which is consistent with the declared rules. In particular, in
`Guess`

, if the handler claims that the guessed letter is in the word
(by passing `True`

to `k`

), there is no way to update the state in
such a way that the number of missing letters or number of guesses does
not follow the rules.

#### Step 4: Implement Interface¶

Having described the rules, and implemented state transitions which
follow those rules as an effect handler, we can now write an interface
for the game which uses the `MYSTERY`

effect:

```
game : Eff () [MYSTERY (Running (S g) w), STDIO]
[MYSTERY NotRunning, STDIO]
```

The type indicates that the game must start in a running state, with some guesses available, and eventually reach a not-running state (i.e. won or lost). The only way to achieve this is by correctly following the stated rules.

Note that the type of `game`

makes no assumption that there are
letters to be guessed in the given word (i.e. it is `w`

rather than
`S w`

). This is because we will be choosing a word at random from a
vector of `String`

, and at no point have we made it explicit that
those `String`

are non-empty.

Finally, we need to initialise the game by picking a word at random from
a list of candidates, setting it as the target using `NewWord`

, then
running `game`

:

```
runGame : Eff () [MYSTERY NotRunning, RND, SYSTEM, STDIO]
runGame = do srand !time
let w = index !(rndFin _) words
call $ NewWord w
game
putStrLn !(call Get)
```

We use the system time (`time`

from the `SYSTEM`

effect; see
Appendix Effects Summary) to initialise the random number
generator, then pick a random `Fin`

to index into a list of
`words`

. For example, we could initialise a word list as follows:

```
words : ?wtype
words = with Vect ["idris","agda","haskell","miranda",
"java","javascript","fortran","basic",
"coffeescript","rust"]
wtype = proof search
```

Note

Rather than have to explicitly declare a type with the vector’s
length, it is convenient to give a hole `?wtype`

and let
Idris’s proof search mechanism find the type. This is a
limited form of type inference, but very useful in practice.

A possible complete implementation of `game`

is
presented below:

```
game : Eff () [MYSTERY (Running (S g) w), STDIO]
[MYSTERY NotRunning, STDIO]
game {w=Z} = Won
game {w=S _}
= do putStrLn !Get
putStr "Enter guess: "
let guess = trim !getStr
case choose (not (guess == "")) of
(Left p) => processGuess (strHead' guess p)
(Right p) => do putStrLn "Invalid input!"
game
where
processGuess : Char -> Eff () [MYSTERY (Running (S g) (S w)), STDIO]
[MYSTERY NotRunning, STDIO]
processGuess {g} {w} c
= case !(Main.Guess c) of
True => do putStrLn "Good guess!"
case w of
Z => Won
(S k) => game
False => do putStrLn "No, sorry"
case g of
Z => Lost
(S k) => game
```

#### Discussion¶

Writing the rules separately as an effect, then an implementation
which uses that effect, ensures that the implementation must follow
the rules. This has practical applications in more serious contexts;
`MysteryRules`

for example can be though of as describing a
*protocol* that a game player most follow, or alternative a
*precisely-typed API*.

In practice, we wouldn’t really expect to write rules first then
implement the game once the rules were complete. Indeed, I didn’t do
so when constructing this example! Rather, I wrote down a set of
likely rules making any assumptions *explicit* in the state
transitions for `MysteryRules`

. Then, when implementing `game`

at
first, any incorrect assumption was caught as a type error. The
following errors were caught during development:

- Not realising that allowing
`NewWord`

to be an arbitrary string would mean that`game`

would have to deal with a zero-length word as a starting state. - Forgetting to check whether a game was won before recursively calling
`processGuess`

, thus accidentally continuing a finished game. - Accidentally checking the number of missing letters, rather than the number of remaining guesses, when checking if a game was lost.

These are, of course, simple errors, but were caught by the type checker before any testing of the game.

[1] | Readers may recognise this game by the name “Hangman”. |

### Further Reading¶

This tutorial has given an introduction to writing and reasoning about
side-effecting programs in Idris, using the `Effects`

library.
More details about the *implementation* of the library, such as how
`run`

works, how handlers are invoked, etc, are given in a separate
paper [1].

Some libraries and programs which use `Effects`

can be found in the
following places:

- https://github.com/edwinb/SDL-idris — some bindings for the SDL media library, supporting graphics in particular.
- https://github.com/edwinb/idris-demos — various demonstration programs, including several examples from this tutorial, and a “Space Invaders” game.
- https://github.com/SimonJF/IdrisNet2 — networking and socket libraries.
- https://github.com/edwinb/Protocols — a high level communication protocol description language.

The inspiration for the `Effects`

library was Bauer and Pretnar’s
Eff language [2], which describes a language based on algebraic
effects and handlers. Other recent languages and libraries have also
been built on this ideas, for example [3] and [4]. The theoretical
foundations are also well-studied see [5], [6], [7], [8].

[1] | Edwin Brady. 2013. Programming and reasoning with algebraic effects and dependent types. SIGPLAN Not. 48, 9 (September 2013), 133-144. DOI=10.1145/2544174.2500581 https://dl.acm.org/citation.cfm?doid=2544174.2500581 |

[2] | Andrej Bauer, Matija Pretnar, Programming with algebraic effects and handlers, Journal of Logical and Algebraic Methods in Programming, Volume 84, Issue 1, January 2015, Pages 108-123, ISSN 2352-2208, http://math.andrej.com/wp-content/uploads/2012/03/eff.pdf |

[3] | Ben Lippmeier. 2009. Witnessing Purity, Constancy and Mutability. In Proceedings of the 7th Asian Symposium on Programming Languages and Systems (APLAS ‘09), Zhenjiang Hu (Ed.). Springer-Verlag, Berlin, Heidelberg, 95-110. DOI=10.1007/978-3-642-10672-9_9 http://link.springer.com/chapter/10.1007%2F978-3-642-10672-9_9 |

[4] | Ohad Kammar, Sam Lindley, and Nicolas Oury. 2013. Handlers in action. SIGPLAN Not. 48, 9 (September 2013), 145-158. DOI=10.1145/2544174.2500590 https://dl.acm.org/citation.cfm?doid=2544174.2500590 |

[5] | Martin Hyland, Gordon Plotkin, John Power, Combining effects: Sum and tensor, Theoretical Computer Science, Volume 357, Issues 1–3, 25 July 2006, Pages 70-99, ISSN 0304-3975, (https://www.sciencedirect.com/science/article/pii/S0304397506002659) |

[6] | Paul Blain Levy. 2004. Call-By-Push-Value: A Functional/Imperative Synthesis (Semantics Structures in Computation, V. 2). Kluwer Academic Publishers, Norwell, MA, USA. |

[7] | Plotkin, Gordon, and Matija Pretnar. “Handlers of algebraic effects.” Programming Languages and Systems. Springer Berlin Heidelberg, 2009. 80-94. |

[8] | Pretnar, Matija. “Logic and handling of algebraic effects.” (2010). |

### Effects Summary¶

This appendix gives interfaces for the core effects provided by the library.

#### EXCEPTION¶

```
module Effect.Exception
import Effects
import System
import Control.IOExcept
EXCEPTION : Type -> EFFECT
raise : a -> Eff b [EXCEPTION a]
Handler (Exception a) Maybe where { ... }
Handler (Exception a) List where { ... }
Handler (Exception a) (Either a) where { ... }
Handler (Exception a) (IOExcept a) where { ... }
Show a => Handler (Exception a) IO where { ... }
```

#### FILE_IO¶

```
module Effect.File
import Effects
import Control.IOExcept
FILE_IO : Type -> EFFECT
data OpenFile : Mode -> Type
open : (fname : String)
-> (m : Mode)
-> Eff Bool [FILE_IO ()]
(\res => [FILE_IO (case res of
True => OpenFile m
False => ())])
close : Eff () [FILE_IO (OpenFile m)] [FILE_IO ()]
readLine : Eff String [FILE_IO (OpenFile Read)]
writeLine : String -> Eff () [FILE_IO (OpenFile Write)]
eof : Eff Bool [FILE_IO (OpenFile Read)]
Handler FileIO IO where { ... }
```

#### RND¶

```
module Effect.Random
import Effects
import Data.Vect
import Data.Fin
RND : EFFECT
srand : Integer -> Eff m () [RND]
rndInt : Integer -> Integer -> Eff m Integer [RND]
rndFin : (k : Nat) -> Eff m (Fin (S k)) [RND]
Handler Random m where { ... }
```

#### SELECT¶

```
module Effect.Select
import Effects
SELECT : EFFECT
select : List a -> Eff m a [SELECT]
Handler Selection Maybe where { ... }
Handler Selection List where { ... }
```

#### STATE¶

```
module Effect.State
import Effects
STATE : Type -> EFFECT
get : Eff m x [STATE x]
put : x -> Eff m () [STATE x]
putM : y -> Eff m () [STATE x] [STATE y]
update : (x -> x) -> Eff m () [STATE x]
Handler State m where { ... }
```

#### STDIO¶

```
module Effect.StdIO
import Effects
import Control.IOExcept
STDIO : EFFECT
putChar : Handler StdIO m => Char -> Eff m () [STDIO]
putStr : Handler StdIO m => String -> Eff m () [STDIO]
putStrLn : Handler StdIO m => String -> Eff m () [STDIO]
getStr : Handler StdIO m => Eff m String [STDIO]
getChar : Handler StdIO m => Eff m Char [STDIO]
Handler StdIO IO where { ... }
Handler StdIO (IOExcept a) where { ... }
```

#### SYSTEM¶

```
module Effect.System
import Effects
import System
import Control.IOExcept
SYSTEM : EFFECT
getArgs : Handler System e => Eff e (List String) [SYSTEM]
time : Handler System e => Eff e Int [SYSTEM]
getEnv : Handler System e => String -> Eff e (Maybe String) [SYSTEM]
Handler System IO where { ... }
Handler System (IOExcept a) where { ... }
```

## Theorem Proving¶

A tutorial on theorem proving in Idris.

Note

The documentation for Idris has been published under the Creative
Commons CC0 License. As such to the extent possible under law, *The
Idris Community* has waived all copyright and related or neighboring
rights to Documentation for Idris.

### Running example: Addition of Natural Numbers¶

Throughout this tutorial, we will be working with the following function, defined in the Idris prelude, which defines addition on natural numbers:

```
plus : Nat -> Nat -> Nat
plus Z m = m
plus (S k) m = S (plus k m)
```

It is defined by the above equations, meaning that we have for free the
properties that adding `m`

to zero always results in `m`

, and that
adding `m`

to any non-zero number `S k`

always results in
`S (plus k m)`

. We can see this by evaluation at the Idris REPL (i.e.
the prompt, the read-eval-print loop):

```
Idris> \m => plus Z m
\m => m : Nat -> Nat
Idris> \k,m => plus (S k) m
\k => \m => S (plus k m) : Nat -> Nat -> Nat
```

Note that unlike many other language REPLs, the Idris REPL performs
evaluation on *open* terms, meaning that it can reduce terms which
appear inside lambda bindings, like those above. Therefore, we can
introduce unknowns `k`

and `m`

as lambda bindings and see how
`plus`

reduces.

The `plus`

function has a number of other useful properties, for
example:

- It is
*commutative*, that is for all`Nat`

inputs`n`

and`m`

, we know that`plus n m = plus m n`

. - It is
*associative*, that is for all`Nat`

inputs`n`

,`m`

and`p`

, we know that`plus n (plus m p) = plus (plus m n) p`

.

We can use these properties in an Idris program, but in order to do so
we must *prove* them.

#### Equality Proofs¶

Idris has a built-in propositional equality type, conceptually defined as follows:

```
data (=) : a -> b -> Type where
Refl : x = x
```

Note that this must be built-in, rather than defined in the library,
because `=`

is a reserved operator — you cannot define this directly
in your own code.

It is *propositional* equality, where the type states that any two
values in different types `a`

and `b`

may be proposed to be equal.
There is only one way to *prove* equality, however, which is by
reflexivity (`Refl`

).

We have a *type* for propositional equality here, and correspondingly a
*program* inhabiting an instance of this type can be seen as a proof of
the corresponding proposition [1]. So, trivially, we can prove that
`4`

equals `4`

:

```
four_eq : 4 = 4
four_eq = Refl
```

However, trying to prove that `4 = 5`

results in failure:

```
four_eq_five : 4 = 5
four_eq_five = Refl
```

The type `4 = 5`

is a perfectly valid type, but is uninhabited, so
when trying to type check this definition, Idris gives the following
error:

```
When elaborating right hand side of four_eq_five:
Type mismatch between
x = x (Type of Refl)
and
4 = 5 (Expected type)
```

##### Type checking equality proofs¶

An important step in type checking Idris programs is *unification*,
which attempts to resolve implicit arguments such as the implicit
argument `x`

in `Refl`

. As far as our understanding of type checking
proofs is concerned, it suffices to know that unifying two terms
involves reducing both to normal form then trying to find an assignment
to implicit arguments which will make those normal forms equal.

When type checking `Refl`

, Idris requires that the type is of the form
`x = x`

, as we see from the type of `Refl`

. In the case of
`four_eq_five`

, Idris will try to unify the expected type `4 = 5`

with the type of `Refl`

, `x = x`

, notice that a solution requires
that `x`

be both `4`

and `5`

, and therefore fail.

Since type checking involves reduction to normal form, we can write the following equalities directly:

```
twoplustwo_eq_four : 2 + 2 = 4
twoplustwo_eq_four = Refl
plus_reduces_Z : (m : Nat) -> plus Z m = m
plus_reduces_Z m = Refl
plus_reduces_Sk : (k, m : Nat) -> plus (S k) m = S (plus k m)
plus_reduces_Sk k m = Refl
```

#### Heterogeneous Equality¶

Equality in Idris is *heterogeneous*, meaning that we can even propose
equalities between values in different types:

```
idris_not_php : 2 = "2"
```

Obviously, in Idris the type `2 = "2"`

is uninhabited, and one might
wonder why it is useful to be able to propose equalities between values
in different types. However, with dependent types, such equalities can
arise naturally. For example, if two vectors are equal, their lengths
must be equal:

```
vect_eq_length : (xs : Vect n a) -> (ys : Vect m a) ->
(xs = ys) -> n = m
```

In the above declaration, `xs`

and `ys`

have different types because
their lengths are different, but we would still like to draw a
conclusion about the lengths if they happen to be equal. We can define
`vect_eq_length`

as follows:

```
vect_eq_length xs xs Refl = Refl
```

By matching on `Refl`

for the third argument, we know that the only
valid value for `ys`

is `xs`

, because they must be equal, and
therefore their types must be equal, so the lengths must be equal.

Alternatively, we can put an underscore for the second `xs`

, since
there is only one value which will type check:

```
vect_eq_length xs _ Refl = Refl
```

#### Properties of `plus`

¶

Using the `(=)`

type, we can now state the properties of `plus`

given above as Idris type declarations:

```
plus_commutes : (n, m : Nat) -> plus n m = plus m n
plus_assoc : (n, m, p : Nat) -> plus n (plus m p) = plus (plus n m) p
```

Both of these properties (and many others) are proved for natural number
addition in the Idris standard library, using `(+)`

from the `Num`

interface rather than using `plus`

directly. They have the names
`plusCommutative`

and `plusAssociative`

respectively.

In the remainder of this tutorial, we will explore several different
ways of proving `plus_commutes`

(or, to put it another way, writing
the function.) We will also discuss how to use such equality proofs, and
see where the need for them arises in practice.

[1] | This is known as the Curry-Howard correspondence. |

### Inductive Proofs¶

Before embarking on proving `plus_commutes`

in Idris itself, let us
consider the overall structure of a proof of some property of natural
numbers. Recall that they are defined recursively, as follows:

```
data Nat : Type where
Z : Nat
S : Nat -> Nat
```

A *total* function over natural numbers must both terminate, and cover
all possible inputs. Idris checks functions for totality by checking that
all inputs are covered, and that all recursive calls are on
*structurally smaller* values (so recursion will always reach a base
case). Recalling `plus`

:

```
plus : Nat -> Nat -> Nat
plus Z m = m
plus (S k) m = S (plus k m)
```

This is total because it covers all possible inputs (the first argument
can only be `Z`

or `S k`

for some `k`

, and the second argument
`m`

covers all possible `Nat`

) and in the recursive call, `k`

is structurally smaller than `S k`

so the first argument will always
reach the base case `Z`

in any sequence of recursive calls.

In some sense, this resembles a mathematical proof by induction (and
this is no coincidence!). For some property `P`

of a natural number
`x`

, we can show that `P`

holds for all `x`

if:

`P`

holds for zero (the base case).- Assuming that
`P`

holds for`k`

, we can show`P`

also holds for`S k`

(the inductive step).

In `plus`

, the property we are trying to show is somewhat trivial (for
all natural numbers `x`

, there is a `Nat`

which need not have any
relation to `x`

). However, it still takes the form of a base case and
an inductive step. In the base case, we show that there is a `Nat`

arising from `plus n m`

when `n = Z`

, and in the inductive step we
show that there is a `Nat`

arising when `n = S k`

and we know we can
get a `Nat`

inductively from `plus k m`

. We could even write a
function capturing all such inductive definitions:

```
nat_induction : (P : Nat -> Type) -> -- Property to show
(P Z) -> -- Base case
((k : Nat) -> P k -> P (S k)) -> -- Inductive step
(x : Nat) -> -- Show for all x
P x
nat_induction P p_Z p_S Z = p_Z
nat_induction P p_Z p_S (S k) = p_S k (nat_induction P p_Z p_S k)
```

Using `nat_induction`

, we can implement an equivalent inductive
version of `plus`

:

```
plus_ind : Nat -> Nat -> Nat
plus_ind n m
= nat_induction (\x => Nat)
m -- Base case, plus_ind Z m
(\k, k_rec => S k_rec) -- Inductive step plus_ind (S k) m
-- where k_rec = plus_ind k m
n
```

To prove that `plus n m = plus m n`

for all natural numbers `n`

and
`m`

, we can also use induction. Either we can fix `m`

and perform
induction on `n`

, or vice versa. We can sketch an outline of a proof;
performing induction on `n`

, we have:

Property

`P`

is`\x => plus x m = plus m x`

.Show that

`P`

holds in the base case and inductive step:- Base case:
`P Z`

, i.e.`plus Z m = plus m Z`

, which reduces to`m = plus m Z`

due to the definition of`plus`

. - Inductive step: Inductively, we know that
`P k`

holds for a specific, fixed`k`

, i.e.`plus k m = plus m k`

(the induction hypothesis). Given this, show`P (S k)`

, i.e.`plus (S k) m = plus m (S k)`

, which reduces to`S (plus k m) = plus m (S k)`

. From the induction hypothesis, we can rewrite this to`S (plus m k) = plus m (S k)`

.

To complete the proof we therefore need to show that `m = plus m Z`

for all natural numbers `m`

, and that `S (plus m k) = plus m (S k)`

for all natural numbers `m`

and `k`

. Each of these can also be
proved by induction, this time on `m`

.

We are now ready to embark on a proof of commutativity of `plus`

formally in Idris.

### Pattern Matching Proofs¶

In this section, we will provide a proof of `plus_commutes`

directly,
by writing a pattern matching definition. We will use interactive
editing features extensively, since it is significantly easier to
produce a proof when the machine can give the types of intermediate
values and construct components of the proof itself. The commands we
will use are summarised below. Where we refer to commands
directly, we will use the Vim version, but these commands have a direct
mapping to Emacs commands.

Command | Vim binding | Emacs binding | Explanation |

Check type | `\t` |
`C-c C-t` |
Show type of identifier or hole under the cursor. |

Proof search | `\o` |
`C-c C-a` |
Attempt to solve hole under the cursor by applying simple proof search. |

Make new definition | `\d` |
`C-c C-s` |
Add a template definition for the type defined under the cursor. |

Make lemma | `\l` |
`C-c C-e` |
Add a top level function with a type which solves the hole under the cursor. |

Split cases | `\c` |
`C-c C-c` |
Create new constructor patterns for each possible case of the variable under the cursor. |

#### Creating a Definition¶

To begin, create a file `pluscomm.idr`

containing the following type
declaration:

```
plus_commutes : (n : Nat) -> (m : Nat) -> n + m = m + n
```

To create a template definition for the proof, press `\d`

(or the
equivalent in your editor of choice) on the line with the type
declaration. You should see:

```
plus_commutes : (n : Nat) -> (m : Nat) -> n + m = m + n
plus_commutes n m = ?plus_commutes_rhs
```

To prove this by induction on `n`

, as we sketched in Section
Inductive Proofs, we begin with a case split on `n`

(press
`\c`

with the cursor over the `n`

in the definition.) You
should see:

```
plus_commutes : (n : Nat) -> (m : Nat) -> n + m = m + n
plus_commutes Z m = ?plus_commutes_rhs_1
plus_commutes (S k) m = ?plus_commutes_rhs_2
```

If we inspect the types of the newly created holes,
`plus_commutes_rhs_1`

and `plus_commutes_rhs_2`

, we see that the
type of each reflects that `n`

has been refined to `Z`

and `S k`

in each respective case. Pressing `\t`

over
`plus_commutes_rhs_1`

shows:

```
m : Nat
--------------------------------------
plus_commutes_rhs_1 : m = plus m 0
```

Note that `Z`

renders as `0`

because the pretty printer renders
natural numbers as integer literals for readability. Similarly, for
`plus_commutes_rhs_2`

:

```
k : Nat
m : Nat
--------------------------------------
plus_commutes_rhs_2 : S (plus k m) = plus m (S k)
```

It is a good idea to give these slightly more meaningful names:

```
plus_commutes : (n : Nat) -> (m : Nat) -> n + m = m + n
plus_commutes Z m = ?plus_commutes_Z
plus_commutes (S k) m = ?plus_commutes_S
```

#### Base Case¶

We can create a separate lemma for the base case interactively, by
pressing `\l`

with the cursor over `plus_commutes_Z`

. This
yields:

```
plus_commutes_Z : m = plus m 0
plus_commutes : (n : Nat) -> (m : Nat) -> n + m = m + n
plus_commutes Z m = plus_commutes_Z
plus_commutes (S k) m = ?plus_commutes_S
```

That is, the hole has been filled with a call to a top level
function `plus_commutes_Z`

. The argument `m`

has been made implicit
because it can be inferred from context when it is applied.

Unfortunately, we cannot prove this lemma directly, since `plus`

is
defined by matching on its *first* argument, and here `plus m 0`

has a
specific value for its *second argument* (in fact, the left hand side of
the equality has been reduced from `plus 0 m`

.) Again, we can prove
this by induction, this time on `m`

.

First, create a template definition with `\d`

:

```
plus_commutes_Z : m = plus m 0
plus_commutes_Z = ?plus_commutes_Z_rhs
```

Since we are going to write this by induction on `m`

, which is
implicit, we will need to bring `m`

into scope manually:

```
plus_commutes_Z : m = plus m 0
plus_commutes_Z {m} = ?plus_commutes_Z_rhs
```

Now, case split on `m`

with `\c`

:

```
plus_commutes_Z : m = plus m 0
plus_commutes_Z {m = Z} = ?plus_commutes_Z_rhs_1
plus_commutes_Z {m = (S k)} = ?plus_commutes_Z_rhs_2
```

Checking the type of `plus_commutes_Z_rhs_1`

shows the following,
which is easily proved by reflection:

```
--------------------------------------
plus_commutes_Z_rhs_1 : 0 = 0
```

For such trivial proofs, we can let write the proof automatically by
pressing `\o`

with the cursor over `plus_commutes_Z_rhs_1`

.
This yields:

```
plus_commutes_Z : m = plus m 0
plus_commutes_Z {m = Z} = Refl
plus_commutes_Z {m = (S k)} = ?plus_commutes_Z_rhs_2
```

For `plus_commutes_Z_rhs_2`

, we are not so lucky:

```
k : Nat
--------------------------------------
plus_commutes_Z_rhs_2 : S k = S (plus k 0)
```

Inductively, we should know that `k = plus k 0`

, and we can get access
to this inductive hypothesis by making a recursive call on k, as
follows:

```
plus_commutes_Z : m = plus m 0
plus_commutes_Z {m = Z} = Refl
plus_commutes_Z {m = (S k)} = let rec = plus_commutes_Z {m=k} in
?plus_commutes_Z_rhs_2
```

For `plus_commutes_Z_rhs_2`

, we now see:

```
k : Nat
rec : k = plus k (fromInteger 0)
--------------------------------------
plus_commutes_Z_rhs_2 : S k = S (plus k 0)
```

Again, the `fromInteger 0`

is merely due to `Nat`

having an implementation
of the `Num`

interface. So we know that `k = plus k 0`

, but how do
we use this to update the goal to `S k = S k`

?

To achieve this, Idris provides a `replace`

function as part of the
prelude:

```
*pluscomm> :t replace
replace : (x = y) -> P x -> P y
```

Given a proof that `x = y`

, and a property `P`

which holds for
`x`

, we can get a proof of the same property for `y`

, because we
know `x`

and `y`

must be the same. In practice, this function can be
a little tricky to use because in general the implicit argument `P`

can be hard to infer by unification, so Idris provides a high level
syntax which calculates the property and applies `replace`

:

```
rewrite prf in expr
```

If we have `prf : x = y`

, and the required type for `expr`

is some
property of `x`

, the `rewrite ... in`

syntax will search for `x`

in the required type of `expr`

and replace it with `y`

. Concretely,
in our example, we can say:

```
plus_commutes_Z {m = (S k)} = let rec = plus_commutes_Z {m=k} in
rewrite rec in ?plus_commutes_Z_rhs_2
```

Checking the type of `plus_commutes_Z_rhs_2`

now gives:

```
k : Nat
rec : k = plus k (fromInteger 0)
_rewrite_rule : plus k 0 = k
--------------------------------------
plus_commutes_Z_rhs_2 : S (plus k 0) = S (plus k 0)
```

Using the rewrite rule `rec`

(which we can see in the context here as
`_rewrite_rule`

[1], the goal type has been updated with `k`

replaced by `plus k 0`

.

Alternatively, we could have applied the rewrite in the other direction
using the `sym`

function:

```
*pluscomm> :t sym
sym : (l = r) -> r = l
```

```
plus_commutes_Z {m = (S k)} = let rec = plus_commutes_Z {m=k} in
rewrite sym rec in ?plus_commutes_Z_rhs_2
```

In this case, inspecting the type of the hole gives:

```
k : Nat
rec : k = plus k (fromInteger 0)
_rewrite_rule : k = plus k 0
--------------------------------------
plus_commutes_Z_rhs_2 : S k = S k
```

Either way, we can use proof search (`\o`

) to complete the
proof, giving:

```
plus_commutes_Z : m = plus m 0
plus_commutes_Z {m = Z} = Refl
plus_commutes_Z {m = (S k)} = let rec = plus_commutes_Z {m=k} in
rewrite rec in Refl
```

The base case is now complete.

#### Inductive Step¶

Our main theorem, `plus_commutes`

should currently be in the following
state:

```
plus_commutes : (n : Nat) -> (m : Nat) -> n + m = m + n
plus_commutes Z m = plus_commutes_Z
plus_commutes (S k) m = ?plus_commutes_S
```

Looking again at the type of `plus_commutes_S`

, we have:

```
k : Nat
m : Nat
--------------------------------------
plus_commutes_S : S (plus k m) = plus m (S k)
```

Conveniently, by induction we can immediately tell that
`plus k m = plus m k`

, so let us rewrite directly by making a
recursive call to `plus_commutes`

. We add this directly, by hand, as
follows:

```
plus_commutes : (n : Nat) -> (m : Nat) -> n + m = m + n
plus_commutes Z m = plus_commutes_Z
plus_commutes (S k) m = rewrite plus_commutes k m in ?plus_commutes_S
```

Checking the type of `plus_commutes_S`

now gives:

```
k : Nat
m : Nat
_rewrite_rule : plus m k = plus k m
--------------------------------------
plus_commutes_S : S (plus m k) = plus m (S k)
```

The good news is that `m`

and `k`

now appear in the correct order.
However, we still have to show that the successor symbol `S`

can be
moved to the front in the right hand side of this equality. This
remaining lemma takes a similar form to the `plus_commutes_Z`

; we
begin by making a new top level lemma with `\l`

. This gives:

```
plus_commutes_S : (k : Nat) -> (m : Nat) -> S (plus m k) = plus m (S k)
```

Unlike the previous case, `k`

and `m`

are not made implicit because
we cannot in general infer arguments to a function from its result.
Again, we make a template definition with `\d`

:

```
plus_commutes_S : (k : Nat) -> (m : Nat) -> S (plus m k) = plus m (S k)
plus_commutes_S k m = ?plus_commutes_S_rhs
```

Again, this is defined by induction over `m`

, since `plus`

is
defined by matching on its first argument. The complete definition is:

```
total
plus_commutes_S : (k : Nat) -> (m : Nat) -> S (plus m k) = plus m (S k)
plus_commutes_S k Z = Refl
plus_commutes_S k (S j) = rewrite plus_commutes_S k j in Refl
```

All holes have now been solved.

The `total`

annotation means that we require the final function to
pass the totality checker; i.e. it will terminate on all possible
well-typed inputs. This is important for proofs, since it provides a
guarantee that the proof is valid in *all* cases, not just those for
which it happens to be well-defined.

Now that `plus_commutes`

has a `total`

annotation, we have completed the
proof of commutativity of addition on natural numbers.

[1] | Note that the left and right hand sides of the equality have been
swapped, because `replace` takes a proof of `x=y` and the
property for `x` , not `y` . |

### DEPRECATED: Interactive Theorem Proving¶

Warning

The interactive theorem-proving interface documented here has been deprecated in favor of Elaborator Reflection.

Idris also supports interactive theorem proving via tactics. This is generally not recommended to be used directly, but rather used as a mechanism for building proof automation which is beyond the scope of this tutorial. In this section, we briefly discus tactics.

One way to write proofs interactively is to write the general *structure* of
the proof, and use the interactive mode to complete the details.
Consider the following definition, proved in Theorem Proving:

```
plusReduces : (n:Nat) -> plus Z n = n
```

We’ll be constructing