Type inference

14 June 2020
This post is part 4 of 4 of the series Introduction to Type Systems.

In the previous post, we have added polymorphism to the simply typed lambda calculus and implemented a type checker for the polymorphic lambda calculus. In this post, we’ll explore type inference or reconstruction.

Imports etc.

module Inference where

import           Control.Monad        (when)
import           Control.Monad.Except
import           Control.Monad.RWS
import           Data.Bifunctor
import qualified Data.Map.Strict      as Map
import           Data.Map.Strict      (Map)
import qualified Data.Set             as Set
import           Data.Set             (Set)
import           Test.QuickCheck


In the polymorphic lambda calculus, we can write polymorphic (generic) functions that work on all types, using parametric polymorphic. This is a major benefit over the simply typed lambda calculus, because it reduces duplication: for example, we no longer have to write an identity function for every type that we might need one for, but can write exactly one identity function that works on all types.

But, as you might have noticed, it is quite some work to use such polymorphic functions. Where we could define \(\mathsf{const}\) as \(\lambda x. \lambda y. x\) and use it like \(\mathsf{const}\ (\lambda x. x)\ (\lambda f. \lambda x. f\ x)\) in the untyped lambda calculus, in the polymorphic lambda calculus we have to type \(\mathsf{const} = \Lambda X. \Lambda Y. \lambda x : X. \lambda y : Y. x\) and use it like the following for the same example: \[ \begin{align*} \mathsf{const}\ & (\forall X. X \rightarrow X) \\ & (\forall A. \forall B. (A \rightarrow B) \rightarrow A \rightarrow B) \\ & (\Lambda X. \lambda x : X. x) \\ & (\Lambda A. \Lambda B. \lambda f : A \rightarrow B. \lambda x : A. f\ x) \end{align*} \]

We have to do a whole lot of typing to make the type checker happy. Wouldn’t it be nice if we could write our terms like in the untyped lambda calculus, with the same static safety as in the polymorphic lambda calculus? It turns out that we can actually implement a type checker that infers or reconstructs the types from a fully untyped program. This technique is called type inference or type reconstruction, and the corresponding type system is called the Hindley-Milner type system.


To write programs without any type information, we remove all types from the syntax of terms. So no more type abstractions, type applications or lambda abstractions with explicit types (e.g., we’ll write \(\lambda x. x\) instead of \(\lambda x : X. x\)).

The AST looks like this:

data Term
  = TmTrue
    -- ^ True value
  | TmFalse
    -- ^ False value
  | TmInt Integer
    -- ^ Integer value
  | TmVar String
    -- ^ Variable
  | TmAbs String Term
    -- ^ Lambda abstraction
  | TmApp Term Term
    -- ^ Application
  | TmAdd Term Term
    -- ^ Addition
  | TmIf Term Term Term
    -- ^ If-then-else conditional
  | TmLet String Term Term
    -- ^ Let-in
  deriving (Show, Eq)

You might notice that this is just the syntax of the untyped lambda calculus (TmVar, TmAbs, TmApp) with the syntax constructs of the simply typed lambda calculus (TmTrue, TmFalse, TmInt, TmAdd, TmIf), plus the addition of the TmLet constructor, which is used for terms of the form \(\mathbf{let}\ x = t\ \mathbf{in}\ t'\). The addition of let-in terms is not strictly necessary, but it is if we actually want to use polymorphism. (This will be discussed later.)

For the syntax of types, we do have to make a substantial change, though. We must restrict our usage of polymorphism: we can only use \(\forall\)’s at the top level; no more \((\forall A. A \rightarrow A) \rightarrow (\forall B. B \rightarrow B)\), for example. We have to do this, because type inference for the polymorphic lambda calculus as we saw it in the previous post is undecidable. We will therefore split our type syntax into two: monotypes and polytypes (or type schemes).

The syntax for polytypes (for which we’ll write \(\sigma\)) is very simple:

\[ \begin{align*} \sigma ::=\ & \forall \vec{X}. \tau & \text{(polytype)} \\ \end{align*} \]

Here \(\tau\) is a monotype, and \(\vec{X}\) is a (possibly empty) list of type variables.

In Haskell, this is:

data Polytype
  = TyForall [String] Type
  deriving (Show, Eq)

(We’ll use just Type to refer to monotypes.)

The syntax for monotypes looks like this:

\[ \begin{align*} \tau ::=\ & X & \text{(type variable)} \\ \mid\ & \tau \rightarrow \tau' & \text{(function type)} \\ \mid\ & \mathsf{Bool} & \text{(boolean type)} \\ \mid\ & \mathsf{Int} & \text{(integer type)} \end{align*} \]

Or in Haskell:

data Type
  = TyVar String
    -- ^ Type variable
  | TyFun Type Type
    -- ^ Function type
  | TyBool
    -- ^ Boolean type
  | TyInt
    -- ^ Integer type
  deriving (Show, Eq)

The type for the identity function (which we now write as just \(\lambda x. x\)), \(\forall X. X \rightarrow X\), is written in Haskell as:

tmId :: Term
tmId = TmAbs "x" (TmVar "x")

tyId :: Polytype
tyId = TyForall ["X"] $ TyFun (TyVar "X") (TyVar "X")

And \(\mathsf{const}\):

tmConst :: Term
tmConst = TmAbs "a" (TmAbs "b" (TmVar "a"))

tyConst :: Polytype
tyConst = TyForall ["A", "B"] $ TyFun (TyVar "A") (TyFun (TyVar "B") (TyVar "A"))

Type checking

Type inference is quite a bit harder than type checking the simply typed lambda calculus or the polymorphic lambda calculus with explicit type annotations. We will use a constraint-based type inference algorithm, based on Types and Programming Languages, Benjamin C. Pierce, Chapter 22.3. I have found this to be the most intuitive approach. I will deviate a bit from Pierce’s approach, though, to make the rules somewhat easier to read.1

For type inference, we will use a different typing relation than the one we used for the simply typed and the polymorphic (but explicitly typed) lambda calculus. Before, we used the relation \(\Gamma \vdash t : \tau\), which could be read something like: \(\Gamma\) entails that \(t\) has type \(\tau\). Now, we will use the typing relation written as follows: \(\Gamma \vdash t : \tau \mid C\). This can be read as: \(\Gamma\) entails that \(t\) has type \(\tau\) if the constraints of \(C\) are satisfied. Our type inference program will generate a set of constraints, which ought to be satisfied for the type checker to succeed. (Another change is the context \(\Gamma\), which will now contain pairs \(x : \sigma\) of variables and polytypes instead of pairs \(x : \tau\) of variables and monotypes.)

Constraint solving

A constraint \(\tau \sim \tau'\) states that \(\tau\) and \(\tau'\) should be unified. The constraint \(A \sim B \rightarrow \mathsf{Int}\), for example, asserts that the type variable \(A\) should be equal to the type \(B \rightarrow \mathsf{Int}\). A constraint set \(C\) is a set (or a list) of constraints. We want to write a a function that unifies a constraint set. This unification function will generate a substitution \(\mathcal{S}\), such that the substitution unifies all constraints in \(C\): for all constraints \(\tau \sim \tau'\), \(\mathcal{S} \tau\) (the substitution \(\mathcal{S}\) applied tot type \(\tau\)) should be equal to \(\mathcal{S} \tau'\).

In Haskell, we will create the following Constraint type, with the infix constructor (:~:) that corresponds to the \(\sim\) in a constraint:

data Constraint = Type :~: Type
  deriving (Show)

For substitutions, we use a map:

type Subst = Map String Type

The substType function will apply a substitution to a type. Applying substitutions to monotypes (i.e., without \(\forall\)s) is quite easy, because we don’t have to worry about renaming.

substType :: Subst -> Type -> Type
substType s TyBool = TyBool
substType s TyInt  = TyInt

When we come across a type variable, we replace it by the corresponding type in the substitution, or keep it when the variable does not occur in the substitution:

substType s (TyVar x) = Map.findWithDefault (TyVar x) x s

For function types, we just apply the substitution recursively:

substType s (TyFun t1 t2) = TyFun (substType s t1) (substType s t2)

With the substType function, we can very easily apply a substitution to a constraint, by applying the substitution to the left-hand side and the right-hand side:

substConstraint :: Subst -> Constraint -> Constraint
substConstraint s (t1 :~: t2) = substType s t1 :~: substType s t2

We can also apply a substitution to a polytype \(\forall \vec{X}. \tau\), which applies the substitution to \(\tau\), with all elements from the substitution with a key from \(\vec{X}\) removed:

substPolytype :: Subst -> Polytype -> Polytype
substPolytype s (TyForall xs ty) =
  let s' = foldr Map.delete s xs
   in TyForall xs (substType s' ty)

As we’ve seen in the previous post, substitution is generally quite hard for types which bind type variables, because the programmer might use the same type variable twice in different contexts, causing them to clash in some cases. Luckily, this won’t be a problem here, since the programmer doesn’t write any type variables. Instead, all type variables that we use are generated by the inference algorithm, which makes sure they are all unique (or fresh). This will be explained later.

We also need to be able to compose two substitutions. In mathematical notation, we write \(\mathcal{S}_1 \circ \mathcal{S}_2\) for the composition of \(\mathcal{S}_1\) and \(\mathcal{S}_2\), where \(\mathcal{S}_2\) is applied first. We want \((\mathcal{S}_1 \circ \mathcal{S}_2)\tau\) for any type \(\tau\) to be equal to \(\mathcal{S}_1(\mathcal{S}_2\tau)\). We first apply \(\mathcal{S}_1\) to the codomain (that is, the values, not the keys, of the Map) of \(\mathcal{S}_2\), and then return the union of the result and \(\mathcal{S}_1\), where values of the first substitution are preferred:

compose :: Subst -> Subst -> Subst
compose s1 s2 = fmap (substType s1) s2 `Map.union` s1

Then, we can write the unification function for a single constraint:

The definition of UnifyError

data UnifyError
  = CannotUnify Type Type
  | InfiniteType String Type
  deriving (Show)
unify :: Constraint -> Either UnifyError Subst
unify c = case c of

To unify two equal simple types, we don’t have to apply any substitution, so we’ll just return an empty substitution:

  TyBool :~: TyBool -> Right Map.empty
  TyInt  :~: TyInt  -> Right Map.empty

To unify two function types, we just need to unify both parameter types and both target types. We do this using the solve function, which can unify a list of constraints. We’ll define solve later.

  TyFun t1 t2 :~: TyFun t1' t2' ->
    solve [ t1 :~: t1'
          , t2 :~: t2' ]

To unify a type variable with another type, we use the bind helper function, which we’ll also define later.

  t1      :~: TyVar x -> bind x t1
  TyVar x :~: t2      -> bind x t2

Any other constraint is unsolvable, so we’ll just throw an error:

  t1 :~: t2 -> Left $ CannotUnify t1 t2

For unifying a type variable with another type, we use the bind function:

bind :: String -> Type -> Either UnifyError Subst
bind x t

When t is the same as the type variable x, we don’t have to do any substituting:

  | t == TyVar x   = Right Map.empty

When the type variable x occurs freely in t (and it is not x itself, which we have checked in the previous case), we cannot unify them, since that would require infinite types. The constraint \(X \sim X \rightarrow X\), for example, has no solution:

  | x `occursIn` t = Left $ InfiniteType x t

Otherwise, we can just return the substitution which substitutes x for t:

  | otherwise      = Right $ Map.fromList [(x, t)]

The occursIn function is very straight-forward:

occursIn :: String -> Type -> Bool
x `occursIn` t = case t of
  TyBool      -> False
  TyInt       -> False
  TyFun t1 t2 -> x `occursIn` t1 || x `occursIn` t2
  TyVar y     -> x == y

Finally, we can solve a list of constraints:

solve :: [Constraint] -> Either UnifyError Subst

Solving an empty list of constraints just corresponds to doing nothing:

solve [] = Right Map.empty

To solve a non-empty list of constraints, we first unify the constraint c, which gives us the substitution s1. We apply this substitution to the rest of the constraints and solve the result, giving us the substitution s2, and then return the composition of s2 and s1:

solve (c:cs) = do
  s1 <- unify c
  s2 <- solve $ fmap (substConstraint s1) cs
  Right (s2 `compose` s1)

Some examples:

solve [TyVar "X" :~: TyInt]
  => Right (fromList [("X",TyBool)])

solve [TyInt :~: TyBool]
  => Left (CannotUnify TyInt TyBool)

solve [TyInt :~: TyVar "X", TyVar "X" :~: TyFun TyBool TyBool]
  => Left (CannotUnify TyInt (TyFun TyBool TyBool))

solve [TyInt :~: TyVar "X", TyVar "Y" :~: TyBool]
  => Right (fromList [("X",TyInt),("Y",TyBool)])

solve [TyVar "X" :~: TyFun (TyVar "X") (TyVar "X")]
  => Left (InfiniteType "X" (TyFun (TyVar "X") (TyVar "X")))

We can also test whether solve has the desired behaviour, namely that the resulting substitution unifies the constraints. To do this, we’ll use the QuickCheck library:

Testing solve

We will first need an instance of Arbitrary for Type and Constraint. The instance for Type is adapted from the lambda calculus example. The frequency for TyInt and TyBool are relatively low, because a frequent occurrence of these simple types in the generated arbitrary types results in a lot of failed calls to solve.

instance Arbitrary Type where
  arbitrary = sized arbType
    arbType n = frequency $
      [ (10, TyVar <$> arbVar)
      , (1, pure TyInt)
      , (1, pure TyBool)
      ] <>
      [ (5, TyFun <$> arbType (n `div` 2) <*> arbType (n `div` 2))
      | n > 0
    arbVar = elements [[c] | c <- ['A'..'Z']]

instance Arbitrary Constraint where
  arbitrary = (:~:) <$> arbitrary <*> arbitrary

Then we write the function unifies, which checks whether a substitution unifies the constraints. (Remember: a substitution \(\mathcal{S}\) satisfies a list of constraints \(C\) if for all constraints \(\tau \sim \tau'\) in \(C\), \(\mathcal{S}\tau = \mathcal{S}\tau'\).)

unifies :: Subst -> [Constraint] -> Bool
unifies s cs =
  let cs' = fmap (substConstraint s) cs
   in all (\(t1 :~: t2) -> t1 == t2) cs'

Now we can write our property, which will check whether every successful solve returns a substitution that unifies the list of constraints. We will discard errors of solve, since they occur quite often for arbitrary constraints, but aren’t useful for checking the property.

prop_solveUnifies :: [Constraint] -> Property
prop_solveUnifies cs =
  case solve cs of
    -- Discard errors
    Left _      -> property Discard
    Right subst -> property $ unifies subst cs

Now we can check the property:

ghci> quickCheck prop_solveUnifies
+++ OK, passed 100 tests; 637 discarded.
Looks good!

Typing rules

Now we know how to solve constraints, but we don’t know how to actually generate them. The typing rules will generate the constraints that should be solves afterwards.

Let’s first look at some easy rules. The rules for the values of the simple types are the still the same as for the simply typed lambda calculus, with the addition of \(\ldots \mid \varnothing\) at the end of the judgement, which states that the rules don’t generate any constraints (an empty set):

The rule for applications is also not that hard:

\[ \text{T-App: } \frac{ \begin{array}{c} \Gamma \vdash t_1 : \tau_1 \mid C_1 \\ \Gamma \vdash t_2 : \tau_2 \mid C_2 \\ X \text{ fresh} \\ C' = C_1 \cup C_2 \cup \{\tau_1 \sim \tau_2 \rightarrow X\} \end{array} }{ \Gamma \vdash t_1\ t_2 : X \mid C' } \]

When type checking the application \(t_1\ t_2\), we first type check \(t_1\) and \(t_2\). We then generate a new constraint set which consists of all the constraints of \(C_1\), all of \(C_2\) and the constraint \(\tau_1 \sim \tau_2 \rightarrow X\). (The \(\cup\) symbol is mathematical notation for the union of two sets.) Because \(t_1\) is applied to \(t_2\), \(t_1\) should be a function with a parameter of the type of \(t_2\). We can’t yet know the resulting type, so we use a fresh type variable, denoted by \(X\), for which we add the constraint that \(\tau_1\) should be equal to \(\tau_2 \rightarrow X\).

To state that \(X\) should be a freshly chosen type variable, we write \(X \text{ fresh}\) in the typing rule. A fresh type variable is a type variable which is not already used elsewhere. Because all terms are implicitly typed (that is, they don’t contain types in their syntax), we can confidently use a predefined list of fresh type variables, since there is no chance of them clashing with type variables written by the programmer (because they don’t exist).

Other rules might add constraints regarding \(X\). The type inference of \(t_1\ t_2 + 3\), for example, will add the constraint \(X \sim \mathsf{Int}\).

The typing rules for if-then-else terms and addition terms are very easy: they are almost the same as for the simply typed lambda calculus, but now we can use constraints to specify that the condition of an if-then-else term must be a boolean, etc.:

\[ \text{T-If: } \frac{ \begin{array}{c} \Gamma \vdash t_1 : \tau_1 \mid C_1 \\ \Gamma \vdash t_2 : \tau_2 \mid C_2 \\ \Gamma \vdash t_3 : \tau_3 \mid C_3 \\ C' = C_1 \cup C_2 \cup C_3 \cup \{\tau_1 \sim \mathsf{Bool}, \tau_2 \sim \tau_3\} \end{array} }{ \Gamma \vdash \mathbf{if}\ t_1\ \mathbf{then}\ t_2\ \mathbf{else}\ t_3 : \tau_2 \mid C' } \]

\[ \text{T-Add: } \frac{ \begin{array}{c} \Gamma \vdash t_1 : \tau_1 \mid C_1 \\ \Gamma \vdash t_2 : \tau_2 \mid C_2 \\ C' = C_1 \cup C_2 \cup \{\tau_1 \sim \mathsf{Int}, \tau_2 \sim \mathsf{Int}\} \end{array} }{ \Gamma \vdash t_1 + t_2 : \mathsf{Int} \mid C' } \]

The rule for variables is a bit more involved. It looks like this:

\[ \text{T-Var: } \frac{ \begin{array}{c} x : \sigma \in \Gamma \\ \tau = \mathit{inst}(\sigma) \end{array} }{ \Gamma \vdash x : \tau \mid \varnothing } \]

Remember that the context \(\Gamma\) contains polytypes, but our typing relation uses monotypes (\(\Gamma \vdash t : \tau\) instead of \(\Gamma \vdash t : \sigma\)). To fix this, we use a function called \(\mathit{inst}\) (short for ‘instantiate’), which takes as its parameter a polytype \(\forall \vec{X}. \tau\). For every type variable \(X_i\) in \(\vec{X}\) (which is a list of type variables), it generates a new, fresh type variable \(Y_i\). It then performs the substitution \([X_1 := Y_1, \ldots, X_n := Y_n]\) on \(\tau\) and returns the result.

This trick is necessary for let-polymorphism (which I’ll discuss in more detail for the typing rule for let-in terms). When inferring the type of the term \[ \begin{array}{l} \mathbf{let}\ \mathsf{id} = \lambda x. x\ \mathbf{in} \\ \mathbf{if}\ \mathsf{id}\ \mathsf{True}\ \mathbf{then}\ \mathsf{id}\ 4\ \mathbf{else}\ 5 \end{array} \] we would add \(\mathsf{id} : \forall A. A \rightarrow A\) to the context. When we come across the term \(\mathsf{id}\ \mathsf{True}\), we would (without using \(\mathit{inst}\)) add the constraint \(A \sim \mathsf{Bool}\). But later, when we type check \(\mathsf{if}\ 4\), we would also add the constraint \(A \sim \mathsf{Int}\). This results in an error, since the unification algorithm can’t unify \(\mathsf{Bool} \sim \mathsf{Int}\) (and rightly so). \(\mathit{inst}\) prevents this problem, as we’ll see when looking at T-Let.

The rule for lambda abstractions looks like this:

\[ \text{T-Abs: } \frac{ \begin{array}{c} X \text{ fresh} \\ \Gamma, x : \forall \varnothing. X \vdash t : \tau \mid C \end{array} }{ \Gamma \vdash \lambda x. t : X \rightarrow \tau \mid C } \]

This can be read as follows: if \(X\) is a free type variable and \(\Gamma, x : \forall \varnothing. X\) entails that \(t\) has type \(\tau\) with the generated constraints \(C\), then \(\Gamma\) entails that \(\lambda x. t\) has type \(X \rightarrow \tau\) with the same generated constraint set \(C\). Since the constraint set stays the same, the T-Abs rule does not introduce any constraints.

Because lambda abstractions are no longer annotated with the type of the parameter (\(\lambda x : \tau. t\)), we don’t know what type we should give \(x\) in the context to type check the body of the lambda abstraction (\(t\)). We therefore use a fresh type variable \(X\) as \(x\)’s type. But, since the context contains polytypes, we can’t just add the pair \(x : X\). We instead add the pair \(x : \forall \varnothing. X\).

Not binding \(X\) with a \(\forall\) (i.e., adding \(x : \forall X. X\)) prevents \(\mathit{inst}\) from applying let-polymorphism to the arguments of lambda abstractions. The above example using a let-in term would not work as a lambda abstraction: \((\lambda \mathsf{id}. \mathbf{if}\ \mathsf{id}\ \mathsf{True}\ \mathbf{then}\ \mathsf{id}\ 4\ \mathbf{else}\ 5)\ (\lambda x. x)\) would fail to type check.

The rule for let-in terms, finally, looks like this:

\[ \text{T-Let: } \frac{ \begin{array}{c} \Gamma \vdash t_1 : \tau_1 \mid C_1 \\ \mathcal{S} = \mathit{solve}(C_1) \\ \sigma = \mathit{gen}(\mathcal{S}\Gamma, \mathcal{S}\tau_1) \\ \Gamma, x : \sigma \vdash t_2 : \tau_2 \mid C_2 \end{array} }{ \Gamma \vdash \mathbf{let}\ x = t_1\ \mathbf{in}\ t_2 : \tau_2 \mid C_2 } \]

This rule is executed in the following steps:

  1. The type of \(t_1\) is determined.
  2. The constraints generated while inferring the type of \(t_1\) are solved using the solve function, giving us the substitution \(\mathcal{S}\).
  3. The substitution is applied to the context \(\Gamma\) and \(\tau_1\) and the resulting type is generalised (using the \(\mathit{gen}\) function). The \(\mathit{gen}\) function creates a polytype \(\sigma\) of the form \(\forall \vec{X}. \mathcal{S}\tau_1\) for the monotype \(\mathcal{S}\tau_1\) in which all free type variables \(\vec{X}\) of \(\mathcal{S}\tau_1\) (not occurring in \(\mathcal{S}\Gamma\)) are bound by a \(\forall\).
  4. The type of \(t_2\) is determined with \(x : \sigma\) added to the context.

This rule adds let-polymorphism to the language. These quite complicated steps are necessary to actually make use of polymorphism. As we saw before, we want lambda abstractions to not support polymorphism, so a parameter can only be used on one concrete type. But for let-in terms, we do want to be able to use the bound variable on multiple concrete types: the identity function on booleans, integers, integer-to-boolean functions, etc.

In the rule for variables, T-Var, we introduced the \(\mathit{inst}\) function. It creates a fresh type variable for every type variable bound in a polytype. To prevent it from generalising the parameters of lambda abstractions, we didn’t bind any type variables in the polytype we added to the context: \(\forall \varnothing. X\). For let-in terms, however, we do want \(\mathit{inst}\) to create another instance for the bound variable for every occurrence. Therefore, we find the most general type for the variable, and add it to the context. When type checking the term \(\mathbf{let}\ \mathsf{id} = \lambda x. x\ \mathbf{in}\ \mathsf{id}\ 1\), for example, \(\mathsf{id}\) is added to the context with its most general type: \(\forall X. X \rightarrow X\). When typing the body of the let-in term, then, the type of \(\mathsf{id}\) is instantiated as \(Y \rightarrow Y\) for example. Then the constraint \(Y \sim \mathsf{Int}\) is generated, because \(\mathsf{id}\) is applied to \(1\), but \(X\) is still untouched.

With these typing rules, we can move on to implementing the type inference algorithm.


For the implementation, we will use so-called monad transformers. However, you should not need to understand how monad transformers work in order to understand the implementation.

Our inference monad looks like this:

type Infer a = RWST Context [Constraint] [String] (Except TypeError) a

type Context = Map String Polytype

The definition of TypeError

data TypeError

The variable was not bound by a lambda abstraction.

  = UnboundVariable String

An error occurred during unification.

  | UnifyError UnifyError
  deriving (Show)

The inference monad is a Reader monad for the Context, which is practically the same as having a Context parameter for every function inside the Infer monad, which is what we did before. Everywhere inside the Infer monad we can get the context, but we can’t change it. Infer is also a Writer for a list of Constraints, which means that we can write to a list of constraints. This list of constraints is the \(\ldots \mid C\) in the typing rules. Infer is furthermore a State for a list of Strings, which will be the supply of fresh type variables. And lastly, Infer can throw TypeErrors.

Using runInfer, we can convert a value of Infer a to an Either TypeError (a, [String], [Constraint]):

runInfer :: Context
         -> [String]
         -> Infer a
         -> Either TypeError (a, [String], [Constraint])
runInfer ctx fs m = runExcept $ runRWST m ctx fs

First, we need a function that generates a fresh type variable. The state should be an infinite list of type variable names, so we should always be able to get the following element from the list:

fresh :: Infer String
fresh = do
  freshVars <- get
  case freshVars of
    []     -> error "Non-infinite list of fresh type variables."
    (f:fs) -> do
      put fs
      pure f

With get :: Infer [String] we can get the list of type variables. When it’s empty, we just use error since the programmer has made a mistake by not using an infinite list of fresh type variables. When the list is non-empty, we return the head, and we use the tail as the new state by using put :: [String] -> Infer (), which replaces the state.

For the initial state of fresh variables, we will use the following:

freshVariables :: [String]
freshVariables = concatMap (\n -> [l : n | l <- ['A'..'Z']]) $
  "" : fmap show [1..]

This list will look something like:

["A", "B", ..., "Z", "A1", "B1", ..., "Z1", "A2", "B2", ...]

We will also need the inst function:

inst :: Polytype -> Infer Type
inst (TyForall xs ty) = do
  ys <- mapM (const fresh) xs
  let subst = Map.fromList $ zip xs (fmap TyVar ys)
  pure $ substType subst ty

For every type variable \(X\) bound by the \(\forall\), we create a fresh type variable \(Y\). Then we apply the substitution which substitutes every \(X_i\) for \(Y_i\).

We also need the gen function, but before we can write it, we need to be able to get the set of free type variables from a type:

freeVarsType :: Type -> Set String
freeVarsType TyBool        = Set.empty
freeVarsType TyInt         = Set.empty
freeVarsType (TyVar x)     = Set.singleton x
freeVarsType (TyFun t1 t2) = freeVarsType t1 `Set.union` freeVarsType t2

And the free type variables from a polytype, which are the free type variables in the monotype that are not bound by the \(\forall\).

freeVarsPolytype :: Polytype -> Set String
freeVarsPolytype (TyForall xs ty) = freeVarsType ty `Set.difference` Set.fromList xs

And also from the context, which corresponds to the union of the free type variables of all polytypes in the context:

freeVarsContext :: Context -> Set String
freeVarsContext = foldMap freeVarsPolytype

Now we can write gen. We will write it outside the Infer monad, because it will be useful elsewhere too.

gen :: Context -> Type -> Polytype
gen ctx ty =
  let xs = Set.toList (freeVarsType ty `Set.difference` freeVarsContext ctx)
   in TyForall xs ty

gen just finds the free type variables of ty which don’t occur in the context, and returns a polytype in which those type variables are bound.

We will also need to be able to apply a substitution to a context, by applying the substitution to every polytype in the context:

substContext :: Subst -> Context -> Context
substContext s = fmap (substPolytype s)

Now we can finally implement the type inference algorithm:

infer :: Term -> Infer Type

\[ \text{T-False: } \frac{ }{ \varnothing \vdash \mathsf{True} : \mathsf{Bool} \mid \varnothing } \]

\[ \text{T-True: } \frac{ }{ \varnothing \vdash \mathsf{True} : \mathsf{Bool} \mid \varnothing } \]

\[ \text{T-Int: } \frac{ }{ \varnothing \vdash n : \mathsf{Int} \mid \varnothing } \]

Values of the simple types are, of course, easy:

infer TmFalse   = pure TyBool
infer TmTrue    = pure TyBool
infer (TmInt _) = pure TyInt

\[ \text{T-App: } \frac{ \begin{array}{c} \Gamma \vdash t_1 : \tau_1 \mid C_1 \\ \Gamma \vdash t_2 : \tau_2 \mid C_2 \\ X \text{ fresh} \\ C' = C_1 \cup C_2 \cup \{\tau_1 \sim \tau_2 \rightarrow X\} \end{array} }{ \Gamma \vdash t_1\ t_2 : X \mid C' } \]

For applications:

infer (TmApp t1 t2) = do

We first infer the types of t1 and t2:

  ty1 <- infer t1
  ty2 <- infer t2

We generate a fresh type variable f:

  f <- TyVar <$> fresh

We generate the constraint ty1 :~: TyFun ty2 f. We can add it to the list of constraints using the tell :: [Constraint] -> Infer () function.

  tell [ty1 :~: TyFun ty2 f]

Finally, we return the fresh type variable as the type:

  pure f

\[ \text{T-If: } \frac{ \begin{array}{c} \Gamma \vdash t_1 : \tau_1 \mid C_1 \\ \Gamma \vdash t_2 : \tau_2 \mid C_2 \\ \Gamma \vdash t_3 : \tau_3 \mid C_3 \\ C' = C_1 \cup C_2 \cup C_3 \cup \{\tau_1 \sim \mathsf{Bool}, \tau_2 \sim \tau_3\} \end{array} }{ \Gamma \vdash \mathbf{if}\ t_1\ \mathbf{then}\ t_2\ \mathbf{else}\ t_3 : \tau_2 \mid C' } \]

For if-then-else terms, we generate the constraints that the condition should be a boolean and that the arms should be of the same type :

infer (TmIf t1 t2 t3) = do
  ty1 <- infer t1
  ty2 <- infer t2
  ty3 <- infer t3
  tell [ty1 :~: TyBool, ty2 :~: ty3]
  pure ty2

\[ \text{T-Add: } \frac{ \begin{array}{c} \Gamma \vdash t_1 : \tau_1 \mid C_1 \\ \Gamma \vdash t_2 : \tau_2 \mid C_2 \\ C' = C_1 \cup C_2 \cup \{\tau_1 \sim \mathsf{Int}, \tau_2 \sim \mathsf{Int}\} \end{array} }{ \Gamma \vdash t_1 + t_2 : \mathsf{Int} \mid C' } \]

The operands of an addition should be integers, and the result is also an integer:

infer (TmAdd t1 t2) = do
  ty1 <- infer t1
  ty2 <- infer t2
  tell [ty1 :~: TyInt, ty2 :~: TyInt]
  pure TyInt

\[ \text{T-Var: } \frac{ \begin{array}{c} x : \sigma \in \Gamma \\ \tau = \mathit{inst}(\sigma) \end{array} }{ \Gamma \vdash x : \tau \mid \varnothing } \]

For variables, we use the inst function:

infer (TmVar x) = do

We can get the context using ask:

  ctx <- ask

We look up x:

  case Map.lookup x ctx of

When it doesn’t exist in the context, we use throwError :: TypeError -> Infer () to throw an error:

    Nothing -> throwError $ UnboundVariable x

Otherwise, we use inst on the type:

    Just ty -> inst ty

\[ \text{T-Abs: } \frac{ \begin{array}{c} X \text{ fresh} \\ \Gamma, x : \forall \varnothing. X \vdash t : \tau \mid C \end{array} }{ \Gamma \vdash \lambda x. t : X \rightarrow \tau \mid C } \]

Then lambda abstractions. Using local :: (Context -> Context) -> Infer a -> Infer a we can update the context for a local sub-computation. To infer the type of t, we need to add x’s type to the context, so we use local. Note that the context is not changed in the outer computation:

infer (TmAbs x t) = do
  f <- TyVar <$> fresh 
  ty <- local (Map.insert x (TyForall [] f)) $ infer t
  pure $ TyFun f ty

\[ \text{T-Let: } \frac{ \begin{array}{c} \Gamma \vdash t_1 : \tau_1 \mid C_1 \\ \mathcal{S} = \mathit{solve}(C_1) \\ \sigma = \mathit{gen}(\mathcal{S}\Gamma, \mathcal{S}\tau_1) \\ \Gamma, x : \sigma \vdash t_2 : \tau_2 \mid C_2 \end{array} }{ \Gamma \vdash \mathbf{let}\ x = t_1\ \mathbf{in}\ t_2 : \tau_2 \mid C_2 } \]

And, finally, let-in terms:

infer (TmLet x t1 t2) = do

We first get the context:

  ctx <- ask

Then we use listen :: Infer a -> Infer (a, [Constraint]) to ‘listen’ to the constraints generated by infer t1. These constraints will not be added to the final list of constraints, but are only generated ‘locally’:

  (ty1, cs) <- listen $ infer t1

Now we try to solve the constraints. If they’re not solvable, we just throw an error. Otherwise, we obtain a substitution:

  subst <- case solve cs of
    Left e  -> throwError $ UnifyError e
    Right s -> pure s

We apply the substitution to t1’s type, ty1, giving us ty1'.

  let ty1' = substType subst ty1

And we generalise ty1' in the context to which we have also applied the substitution, giving us a polytype s:

  let s = gen (substContext subst ctx) ty1'

We add s to the context and infer t2’s type:

  local (Map.insert x s) $ infer t2

That’s it! We’ve written an function which runs the inference algorithm on a term, giving us a type and a list of constraints.

Now, we still need to solve the constraints and apply the substitution to the type. We will write the function polytypeOf, which runs the inference algorithm, solves the constraints, applies the substitution, and turns the resulting type into a polytype:

polytypeOf :: Term -> Either TypeError Polytype
polytypeOf t = do

Run the inference algorithm in an empty context2, giving us a type ty, a list of fresh variables fs and a list of constraints cs:

  (ty, fs, cs) <- runInfer Map.empty freshVariables $ infer t

Solve the constraints to obtain a substitution. Because solve returns an Either UnifyError Subst, we need to turn its error into a TypeError, which we can do by applying the type constructor TypeError to it. To do this, we use first :: Bifunctor p => (a -> b) -> p a c -> p b c:

  subst <- first UnifyError $ solve cs

We apply the substitution to ty:

  let ty' = substType subst ty

We generalise the type in an empty context, giving us the polytype s:

  let s = gen Map.empty ty'

And we return s:

  Right s

Let’s try it!

The type of id:

polytypeOf tmId
  => Right (TyForall ["A"] (TyFun (TyVar "A") (TyVar "A")))

That is \(\forall A. A \rightarrow A\), correct!

The type of const:

polytypeOf tmConst
  => Right (TyForall ["A","B"] (TyFun (TyVar "A") (TyFun (TyVar "B") (TyVar "A"))))

\(\forall A\ B. A \rightarrow B \rightarrow A\), again correct!

Now let’s try to use let-polymorphism, by trying the term: \[ \begin{array}{l} \mathbf{let}\ \mathsf{id} = \lambda x. x\ \mathbf{in} \\ \mathbf{if}\ \mathsf{id}\ \mathsf{True}\ \mathbf{then}\ \mathsf{id}\ 4\ \mathbf{else}\ 5 \end{array} \]

polytypeOf (TmLet "id" (TmAbs "x" (TmVar "x")) (TmIf (TmApp (TmVar "id") TmTrue) (TmApp (TmVar "id") (TmInt 4)) (TmInt 5)))
  => Right (TyForall [] TyInt)

And the same term, but using a lambda abstraction:

\[ (\lambda \mathsf{id}. \mathbf{if}\ \mathsf{id}\ \mathsf{True}\ \mathbf{then}\ \mathsf{id}\ 4\ \mathbf{else}\ 5)\ (\lambda x. x) \]

polytypeOf (TmApp (TmAbs "id" (TmIf (TmApp (TmVar "id") TmTrue) (TmApp (TmVar "id") (TmInt 4)) (TmInt 5))) (TmAbs "x" (TmVar "x")))
  => Left (UnifyError (CannotUnify TyBool TyInt))

Just like we expected, it can’t unify \(\mathsf{Bool} \sim \mathsf{Int}\).

One more: \[ \begin{array}{l} \mathbf{let}\ \mathsf{id} = \lambda x. x\ \mathbf{in} \\ \mathbf{let}\ \mathsf{const} = \lambda a. \lambda b. a\ \mathbf{in} \\ \mathsf{const}\ \mathsf{id}\ \mathsf{const} \end{array} \]

polytypeOf $ TmLet "id" tmId $ TmLet "const" tmConst $ TmApp (TmApp (TmVar "const") (TmVar "id")) (TmVar "const")
  => Right (TyForall ["F"] (TyFun (TyVar "F") (TyVar "F")))

It returns \(\forall F. F \rightarrow F\), which is exactly the type of \(\mathsf{id}\).


We’ve explored Hindley-Milner type inference, and implemented a type inference algorithm! This language is already quite close to Haskell.

Some exercises you might like to do:

  1. Write a function simplPolytype which ‘simplifies’ a polytype. It should rename the bound variables in a polytype to names in the beginning of the alphabet (or: the beginning of freshVariables). The polytype of the last example is \(\forall F. F \rightarrow F\), for example, but it would be nicer if polytypeOf returned \(\forall A. A \rightarrow A\).
  2. Extend the language using other simple types and operations for them.

And, if you have trouble understanding some parts, try to experiment with them a lot. And feel free to ask questions on Reddit.

Further reading

Other resources you might find useful:

  1. Pierce uses the typing relation \(\Gamma \vdash t : \tau \mid_X C\), where the set \(X\) keeps track of the used type variables. This is very useful to formally reason about the type inference algorithm, but it makes the typing rules more complex than necessary for a Haskell implementation. Instead, I will just write \(X \text{ fresh}\) for a type variable \(X\). This approach is more informal, since it doesn’t formally specify when a variable is fresh, but I think it is easier.↩︎

  2. If you want to extend the language by having declarations, or by making a REPL, you might want to run infer in a specific context, so declarations aren’t lost. You would also have to run gen with this context, instead of an empty context.↩︎