CSE 230 W11 - Type Inference

Adapted from the superb notes of Martin Grabmueller

> {-# LANGUAGE FlexibleContexts, TypeSynonymInstances, NoMonomorphismRestriction, OverlappingInstances, FlexibleInstances #-}
> import qualified Data.Map as Map
> import qualified Data.Set as Set
> import Data.Functor ((<$>))
> import Control.Monad.Error
> import Control.Monad.Reader
> import Control.Monad.State
> import qualified Text.PrettyPrint as PP

In this lecture, we will formalize the Hindley-Milner type inference algorithm, a function of whose type is

> typeInference :: TypeEnv -> Exp -> Either String Type

That is, the function takes as input a type-environment containing the types for variables (think “symbol table”), an expression whose type is to be inferred, and returns either an error string, if the expression has a type error, or the type inferred for the expression.


We start by defining the datatypes (abstract syntax) for


First, we define expressions, which is a direct translation of the different cases of lambda-calculus expressions.

> data Exp     =  EVar EVar 
> | ELit Lit
> | EApp Exp Exp
> | EAbs EVar Exp
> | ELet EVar Exp Exp
> deriving (Eq, Ord)

We use a wrapped variant of String to represent program variables

> newtype EVar = EV String deriving (Eq, Ord)

and we use Lit to denote the primitive constants in the language

> data Lit     =  LInt Integer
> | LBool Bool
> deriving (Eq, Ord)

Types and Schemes

Next, we define the different types for our language. In essence, a type is either a base type TInt or TBool, or function types denoted by TArr or a type variables TVar.

> data Type    =  TVar TVar 
> | TInt
> | TBool
> | Type `TArr` Type
> deriving (Eq, Ord)
> newtype TVar = TV String deriving (Eq, Ord)
> data Scheme = Forall [TVar] Type

In order to provide readable output and error messages, we define several pretty-printing functions for the abstract syntax, that are shown at the end of this file.


The last datatype needed to formalize the behavior of the type inference algorithm is an environment (think “symbol table”) that contains the type schemes of all the free variables that appear in the expression to be analyzed. In other words, the environment maps each externally defined variable to its scheme

> newtype TypeEnv = TypeEnv (Map.Map EVar Scheme)

We will use the helper function env \\ (x, s) to replace the binding for variable x in the environment env with s

> (\\) :: TypeEnv -> (EVar, Scheme) -> TypeEnv
> (TypeEnv env) \\ (x, s) = TypeEnv $ Map.insert x s env


As we saw in our informal overview, the algorithm proceeds by generating fresh type variables for unknown types, and then traverses the expression to equate or unify the types of the sub-expressions based on how they are used with each other. This unification, in turn, proceeds by substituting an occurrence of a type variable’s occurrence in a type, with a whole type (eg substituting a type variable with some function type.) Next, we formalize the notion of substitution, and use it to define the unification procedure.

Free Variables and Substitution

First, we define substitutions, which are finite mappings from type variables to types.

> type Subst = Map.Map TVar Type

Next, we define the function apply that takes a substitution and applies it to a type by replacing variables in the type with the corresponding mapping in the substitution (if one exists.)

For example, if we have the substitution that maps type variable a to the type c -> d, then when we apply the substitution to the type a -> a we should get the output (c -> d) -> (c -> d).

But lookout! What if we were to apply the substitution to the scheme Forall a. a -> a -> b ? Now, due to the quantifier, the scheme is α-equivalent (ie the same except for the name of the quantified variable) to Forall z. z -> z -> b.

Thus, we should not substitute the a inside the scheme as they are bound by the quantifier. Instead, we should only substitute instances of free type variables, like b which are not bound by any quantifier. And so, we need to be careful about the free type variables, when applying substitutions.

It turns out that we will have to define the apply function for other kinds of values, like schemes, environments and so on, and so it is useful to create a typeclass

> class Substitutable a where 
> apply :: Subst -> a -> a
> freeTvars :: a -> Set.Set TVar

that contains two operations, which respectively carry out the substitution, and determine the set of free type variables.

Next, we formally define substitutions for types

> instance Substitutable Type where
> apply _ TInt = TInt
> apply _ TBool = TBool
> apply su t@(TVar a) = Map.findWithDefault t a su
> apply su (t1 `TArr` t2) = apply su t1 `TArr` apply su t2
>   freeTvars TInt           =  Set.empty
> freeTvars TBool = Set.empty
> freeTvars (TVar a) = Set.singleton a
> freeTvars (t1 `TArr` t2) = freeTvars t1 `Set.union` freeTvars t2

and schemes

> instance Substitutable Scheme where
> apply s (Forall as t) = Forall as $ apply s' t
> where s' = foldr Map.delete s as
> freeTvars (Forall as t) = (freeTvars t) `Set.difference` (Set.fromList as)

Note that the free variables in a scheme are those in the underlying type minus the variables quantified over by the scheme (since those are bound inside the scheme.) Further, the apply function avoids substituting the bound variables, by using a substitution s' from which the mappings for the scheme’s quantified variables have been deleted.

Often, we will want to lift Substitutable to lists

> instance Substitutable a => Substitutable [a] where
> apply = map . apply
> freeTvars = foldr Set.union Set.empty . map freeTvars

We use the above to make environments Substitutable

> instance Substitutable TypeEnv where
> apply s (TypeEnv env) = TypeEnv $ Map.map (apply s) env
> freeTvars (TypeEnv env) = freeTvars $ Map.elems env

It is useful to define an empty substitution (that leaves a type unchanged)

> empSubst  ::  Subst
> empSubst = Map.empty

and we can compose two substitutions as su1 after su2

> after         :: Subst -> Subst -> Subst
> su1 `after` su2 = (Map.map (apply su1) su2) `Map.union` su1

which yields a single substitution that carries out the substitutions in su2 after which it carries out those in su1.

Most General Unifier

Armed with the above, we can now formally define the notion of type unification. Recall that we want unification to have the following example informal behavior:

a `Int`Inta := Int
ab `b`a := b
a->ba->da->d `b := d`
a->IntBool->bBool->Inta := Bool, b:=Int

The first few cases are where unification is indeed possible, and the last few cases are where it fails corresponding to a type error in the source program. The very last case is an interesting one; the failure is because the type variable a in the first type occurs free inside the second type. Thus, if we try substituting a with a->Int we will just keep spinning forever! Hence, this also throws a unification failure.

Exercise: Write a Haskell program that is rejected by the typechecker because it fails the above occurs check. This is not difficult, the chances are you’ve done this any number of times already!

Here is the unification function mgu that takes two types as input and either returns a successful unified output along with the substitution (as shown in the table above) or an error string explaining the failure (hence, our use of an error monad to describe output.)

> mgu (l `TArr` r) (l' `TArr` r')  = do  s1 <- mgu l l'
> s2 <- mgu (apply s1 r) (apply s1 r')
> return (s1 `after` s2)
> mgu (TVar a) t = varAsgn a t
> mgu t (TVar a) = varAsgn a t
> mgu TInt TInt = return empSubst
> mgu TBool TBool = return empSubst
> mgu t1 t2 = throwError $ "types do not unify: " ++ show t1 ++ " vs. " ++ show t2

The function varAsgn attempts to assign a type variable to a type and return that assignment as a subsitution, but throws an error if the variable occurs within the assigned type.

> varAsgn a t 
> | t == TVar a = return empSubst
> | a `Set.member` (freeTvars t) = throwError $ "occur check fails: " ++ show a ++ " in " ++ show t
> | otherwise = return $ Map.singleton a t

The name mgu stands for Most-General-Unifier ; the function is guaranteed to find the most general unification possible (that is not unify a and b to Int via the substitution a := Int, b:= Int.) This property is crucial for showing that type inference returns the most general type possible for any term (that is, Forall a. a -> a and not Int -> Int for the identity function \x -> x)

The Type Inference Algorithm

We are almost ready to put the pieces together and see how type inference works.

Generalization and Instantiation

The two remaining parts of the puzzle are the operations

Here is the code for the generalize function which simply abstracts a type over all type variables which are free in the type but are not free (and hence, unconstrained) in the given type environment.

> generalize        ::  TypeEnv -> Type -> Scheme
> generalize env t = Forall as t
> where as = Set.toList $ (freeTvars t) `Set.difference` (freeTvars env)

The instantiation function replaces all bound type variables in a type scheme with fresh (and hence, unconstrained) type variables. Now this slightly tricky; how do we generate a fresh variable that is distinct from all previously generated variables?

We need some state to know what type variables have already generated, and so we will use a MonadState with the following type representing the number of previously generated variables. (Aside: This is exactly like the mlabel function from here)

> data TIState = TIState { count :: Int }

Now, the following action yields a fresh Int

> fresh = do s     <- get
> let n = count s
> put $ s { count = n + 1 }
> return n

and we use it to build a function that generates a new type variable with a given prefix

> freshTVar prefix = fresh >>= return . TVar . TV . (prefix ++) . show

Now that we can generate fresh type variables, we can define the instantiate function as

> instantiate (Forall as t) = do as' <- mapM (\ _ -> freshTVar "a") as 
> let s = Map.fromList $ zip as as'
> return $ apply s t

Several operations, for example type scheme instantiation, require fresh names for newly introduced type variables. This is implemented by using an appropriate monad which takes care of generating fresh names. It is also capable of passing a dynamically scoped environment, error handling and performing I/O, but we will not go into details here.

Main Type Inference Function

Now, we are ready to pin down the type inference function ti which infers the types for expressions.

> ti ::  (MonadState TIState m, MonadError String m) => 
> TypeEnv -> Exp -> m (Subst, Type)

The function expects the precondition that the type environment must contain bindings for all free variables of the expressions. The output action is a state-and-error monad m containing a pair of a substitution which records the type constraints imposed on type variables by the expression, and the inferred type of the expression.

First, the easiest cases are the literals whose base types are trivially inferred.

> ti env (ELit (LInt _))  = return (empSubst, TInt)
> ti env (ELit (LBool _)) = return (empSubst, TInt)

Next, consider the case of variables. Here we simply lookup the environment for the scheme for the variable, and if found, create a fresh instantiation for this particular usage of the variable.

> ti (TypeEnv env) (EVar x) = 
> case Map.lookup x env of
> Nothing -> throwError $ "unbound variable: " ++ show x
> Just s -> instantiate s >>= return . (,) empSubst

For the EAbs x e case (ie function definition) we assign the formal x a fresh type variable and analyze the body e using that type variable. Note how the substitution from the body is applied to the type variable for the formal to obtain the inferred function type.

> ti env (EAbs x e) =
> do tv <- freshTVar "a"
> let env' = env \\ (x, Forall [] tv)
> (s1, t1) <- ti env' e
> return (s1, (apply s1 tv) `TArr` t1)

Next, notice how in the EApp e1 e2 case the recorded substitutions for e1 are applied to the environment before e2 is analyzed, and the substitutions from e2 are composed with those for e1 to yield the substitution for the entire expression.

> ti env (EApp e1 e2) =
> do tv <- freshTVar "a"
> (s1, t1) <- ti env e1
> (s2, t2) <- ti (apply s1 env) e2
> s3 <- mgu (apply s2 t1) (TArr t2 tv)
> return (s3 `after` s2 `after` s1, apply s3 tv)

Finally, for the case of a let-binding ELet x e1 e2 we infer the type for e1 and then generalize it to obtain the scheme bound to x when analyzing e2.

> ti env (ELet x e1 e2) =
> do (s1, t1) <- ti env e1
> let t' = generalize (apply s1 env) t1
> env' = env \\ (x, t')
> (s2, t2) <- ti (apply s1 env') e2
> return (s1 `after` s2, t2)

Finally, we use the ti function (which returns a type inference action) to define the typeInference function promised at the beginning.

> -- type TI a = ErrorT String (State TIState) a
> typeInference env e = uncurry apply <$> res
> where act = ti env e
> res = evalState (runErrorT act) s0
> s0 = TIState { count = 0 }


Here are some simple expressions that we can use to test typeInference

> e0  =  ELet (EV "id") (EAbs (EV "x") (EVar (EV "x")))
> (EVar (EV "id"))
> e1 = ELet (EV "id") (EAbs (EV "x") (EVar (EV "x")))
> (EApp (EVar (EV "id")) (EVar (EV "id")))
> e2 = ELet (EV "id") (EAbs (EV "x") (ELet (EV "y") (EVar (EV "x")) (EVar (EV "y"))))
> (EApp (EVar (EV "id")) (EVar (EV "id")))
> e3 = ELet (EV "id") (EAbs (EV "x") (ELet (EV "y") (EVar (EV "x")) (EVar (EV "y"))))
> (EApp (EApp (EVar (EV "id")) (EVar (EV "id"))) (ELit (LInt 2)))
> e4 = ELet (EV "id") (EAbs (EV "x") (EApp (EVar (EV "x")) (EVar (EV "x"))))
> (EVar (EV "id"))
> e5 = EAbs (EV "m") (ELet (EV "y") (EVar (EV "m"))
> (ELet (EV "x") (EApp (EVar (EV "y")) (ELit (LBool True)))
> (EVar (EV "x"))))

This simple test function tries to infer the type for the given expression. If successful, it prints the expression together with its type, otherwise, it prints the error message.

> test :: Exp -> IO ()
> test e = case typeInference (TypeEnv Map.empty) e of
> Left err -> putStrLn $ "error: " ++ err
> Right t -> putStrLn $ show e ++ " :: " ++ show t

The main program simply infers the types for all the example expressions and prints them together with their inferred types, or prints an error message if type inference fails.

> main :: IO ()
> main = mapM_ test [e0, e1, e2, e3, e4, e5]

Appendix: Pretty-printing

This appendix defines pretty-printing functions and instances for Show for all interesting type definitions.

> instance Show TVar where
> showsPrec _ x = shows (prTVar x)
> prTVar (TV a) = PP.text a
> instance Show Type where
> showsPrec _ x = shows (prType x)
> prType :: Type -> PP.Doc
> prType (TVar a) = prTVar a
> prType TInt = PP.text "Int"
> prType TBool = PP.text "Bool"
> prType (TArr t s) = prParenType t PP.<+> PP.text "->" PP.<+> prType s
> prParenType :: Type -> PP.Doc
> prParenType t = case t of
> TArr _ _ -> PP.parens (prType t)
> _ -> prType t
> instance Show EVar where
> showsPrec _ x = shows (prEVar x)
> instance Show Exp where
> showsPrec _ x = shows (prExp x)
> prEVar (EV x) = PP.text x
> prExp :: Exp -> PP.Doc
> prExp (EVar x) = prEVar x
> prExp (ELit lit) = prLit lit
> prExp (ELet x b body) = PP.text "let" PP.<+>
> prEVar x PP.<+> PP.text "=" PP.<+>
> prExp b PP.<+> PP.text "in" PP.$
> PP.nest 2 (prExp body)
> prExp (EApp e1 e2) = prExp e1 PP.<+> prParenExp e2
> prExp (EAbs x e) = PP.char '\\' PP.<+> prEVar x PP.<+>
> PP.text "->" PP.<+>
> prExp e
> prParenExp :: Exp -> PP.Doc
> prParenExp t = case t of
> ELet _ _ _ -> PP.parens (prExp t)
> EApp _ _ -> PP.parens (prExp t)
> EAbs _ _ -> PP.parens (prExp t)
> _ -> prExp t
> instance Show Lit where
> showsPrec _ x = shows (prLit x)
> prLit :: Lit -> PP.Doc
> prLit (LInt i) = PP.integer i
> prLit (LBool b) = if b then PP.text "True" else PP.text "False"
> instance Show Scheme where
> showsPrec _ x = shows (prScheme x)
> prScheme :: Scheme -> PP.Doc
> prScheme (Forall as t) = PP.text "All" PP.<+>
> PP.hcat (PP.punctuate PP.comma (map prTVar as))
> PP.<> PP.text "." PP.<+> prType t