I’ve been working on a simple Haskell98 compiler over the last few days, partly as an excuse to learn how it works, and partly to have a test-bed for trying out some potential language extensions. More on that in a future blog post.

As of yesterday, I have typeclass resolution working. The algorithm to desugar constraints into dictionaries hasn’t been discussed much. Since it’s rather involved, and quite interesting, I thought it might make a good topic for a blog post.

Our journey begins having just implemented Algorithm W aka Hindley-Milner. This is pretty well described in the literature, and there exist several implementations of it in Haskell, so we will not dally here. Algorithm W cashes out in a function of the type:

infer :: SymTable VName Type -> Exp VName -> TI Type

where SymTable VName is a mapping from identifiers in scope to their types, Exp VName is an expression we want to infer, and TI is our type-inference monad. As a monad, TI gives us the ability to generate fresh type variables, and to unify types as we go. Type represents an unqualified type, which is to say it can be used to describe the types a , and Int , but not Eq a => a . We will be implementing qualified types in this blog post.

infer is implemented as a catamorphism, which generates a fresh type variable for every node in the expression tree, looks up free variables in the SymTable and attempts to unify as it goes.

The most obvious thing we need to do in order to introduce constraints to our typechecker is to be able to represent them, so we two types:

infixr 0 :=> data Qual t = ( :=> ) { qualPreds :: [ Pred ] , unqualType :: t } deriving ( Eq , Ord , Functor , Traversable , Foldable ) data Pred = IsInst { predCName :: TName , predInst :: Type } deriving ( Eq , Ord )

Cool. A Qual Type is now a qualified type, and we can represent Eq a => a via [IsInst "Eq" "a"] :=> "a" (assuming OverloadedStrings is turned on.) With this out of the way, we’ll update the type of infer so its symbol table is over Qual Types , and make it return a list of Pred s:

infer :: SymTable VName ( Qual Type ) -> Exp VName -> TI ([ Pred ], Type ) ([],

We update the algebra behind our infer catamorphism so that adds any Pred s necessary when instantiating types:

V a) = infer sym (a) case lookupSym a sym of lookupSym a sym Nothing -> throwE $ "unbound variable: '" <> show a <> "'" throwE Just sigma -> do sigma :=> t) <- instantiate a sigma (pst)instantiate a sigma pure (ps, t) (ps, t)

and can patch any other cases which might generate Pred s. At the end of our cata, we’ll have a big list of constraints necessary for the expression to typecheck.

As a first step, we’ll just write the type-checking part necessary to implement this feature. Which is to say, we’ll need a system for discharging constraints at the type-level, without necessarily doing any work towards code generation.

Without the discharging step, for example, our algorithm will typecheck (==) (1 :: Int) as Eq Int => Int -> Bool , rather than Int -> Bool (since it knows Eq Int .)

Discharging is a pretty easy algorithm. For each Pred , see if it matches the instance head of any instances you have in scope; if so, recursively discharge all of the instance’s context. If you are unable to find any matching instances, just keep the Pred . For example, given the instances:

instance Eq Int instance ( Eq a, Eq b) => Eq (a, b) a,b)(a, b)

and a IsInst "Eq" ("Int", "c") , our discharge algorithm will look like this:

discharging: Eq (Int, c) try: Eq Int --> does not match try: Eq (a, b) --> matches remove `Eq (Int, c)` pred match types: a ~ c b ~ Int discharge: Eq Int discharge: Eq c discharging: Eq Int try: Eq Int --> matches remove `Eq Int` pred discharging: Eq c try: Eq Int --> does not match try: Eq (a, b) --> does not match keep `Eq c` pred

We can implement this in Haskell as:

match :: Pred -> Pred -> TI ( Maybe Subst ) getInsts :: ClassEnv -> [ Qual Pred ] discharge :: ClassEnv -> Pred -> TI ( Subst , [ Pred ]) , []) = do discharge cenv p -- find matching instances and return their contexts <- matchingInstances $ \(qs :=> t) -> do for (getInsts cenv)\(qst) -- the alternative here is to prevent emitting kind -- errors if we compare this 'Pred' against a -- differently-kinded instance. <- ( fmap (qs,) <$> match t p) <|> pure Nothing res(qs,)match t p) pure $ First res res case getFirst $ mconcat matchingInstances of getFirstmatchingInstances Just (qs, subst) -> (qs, subst) -- match types in context let qs' = sub subst qs qs'sub subst qs -- discharge context fmap mconcat $ traverse (discharge cenv) qs' (discharge cenv) qs' Nothing -> -- unable to discharge pure ( mempty , pure p) p)

Great! This works as expected, and if we want to only write a type-checker, this is sufficient. However, we don’t want to only write a type-checker, we also want to generate code capable of using these instances too!

We can start by walking through the transformation in Haskell, and then generalizing from there into an actual algorithm. Starting from a class definition:

class Functor f where fmap :: (a -> b) -> f a -> f b (ab)f af b

we will generate a dictionary type for this class:

data @ Functor f = @ Functor { @ fmap :: (a -> b) -> f a -> f b (ab)f af b }

(I’m using the @ signs here because these things are essentially type applications. That being said, there will be no type applications in this post, so the @ should always be understood to be machinery generated by the compiler for dictionary support.)

Such a definition will give us the following terms:

@ Functor :: ((a -> b) -> f a -> f b) -> @ Functor f ((ab)f af b) @ fmap :: @ Functor f -> (a -> b) -> f a -> f b (ab)f af b

Notice that @fmap is just fmap but with an explicit dictionary ( @Functor f ) being passed in place of the Functor f constraint.

From here, in order to actually construct one of these dictionaries, we can simply inline an instances method:

instance Functor Maybe where fmap = \f m -> case m of { Just x -> Just (f x); Nothing -> Nothing } \f m(f x); -- becomes @ Functor @ Maybe :: @ Functor Maybe @ Functor @ Maybe = @ Functor { @ fmap = \f m -> case m of { Just x -> Just (f x); Nothing -> Nothing } fmap\f m(f x); }

Now we need to look at how these dictionaries actually get used. It’s clear that every fmap in our expression tree should be replaced with @fmap d for some d . If the type of d is monomorphic, we can simply substitute the dictionary we have:

x :: Maybe Int x = fmap ( + 5 ) ( Just 10 ) ) ( -- becomes x :: Maybe Int x = @ fmap @ Functor @ Maybe ( + 5 ) ( Just 10 ) fmap) (

but what happens if the type f is polymorphic? There’s no dictionary we can reference statically, so we’ll need to take it as a parameter:

y :: Functor f => f Int -> f Int y = \z -> fmap ( + 5 ) z \z) z -- becomes y :: @ Functor f -> f Int -> f Int y = \d -> \z -> @ fmap d ( + 5 ) z \d\zfmap d () z

A reasonable question is when should we insert these lambdas to bind the dictionaries? This stumped me for a while, but the answer is whenever you get to a binding group; which is to say whenever your expression is bound by a let , or whenever you finish processing a top-level definition.

One potential gotcha is what should happen in the case of instances with their own contexts? For example, instance (Eq a, Eq b) => Eq (a, b) ? Well, the same rules apply; since a and b are polymorphic constraints, we’ll need to parameterize our @Eq@(,) dictionary by the dictionaries witnessing Eq a and Eq b :

instance ( Eq a, Eq b) => Eq (a, b) where a,b)(a, b) ( == ) = \ab1 ab2 -> ( == ) ( fst ab1) ( fst ab2) \ab1 ab2) (ab1) (ab2) && ( == ) ( snd ab1) ( snd ab2) ) (ab1) (ab2) -- becomes @ Eq @ (,) :: @ Eq a -> @ Eq b -> @ Eq (a, b) (,)(a, b) @ Eq @ (,) = \d1 -> \d2 -> (,)\d1\d2 @ Eq @== ) = \ab1 ab2 -> ( @== ) d1 ( fst ab1) ( fst ab2) { (\ab1 ab2) d1 (ab1) (ab2) && ( @== ) d2 ( snd ab1) ( snd ab2) ) d2 (ab1) (ab2) }

Super-class constraints behave similarly.

So with all of the theory under our belts, how do we actually go about implementing this? The path forward isn’t as straight-forward as we might like; while we’re type-checking we need to desugar terms with constraints on them, but the result of that desugaring depends on the eventual type these terms receive.

For example, if we see (==) in our expression tree, we want to replace it with (@==) d where d might be @Eq@Int , or it might be @Eq@(,) d1 d2 , or it might just stay as d ! And the only way we’ll know what’s what is after we’ve performed the dischargement of our constraints.

As usual, the solution is to slap more monads into the mix:

infer :: SymTable VName ( Qual Type ) -> Exp VName -> TI ( [ Pred ] ( [ , Type , Reader ( Pred -> Exp VName ) ( Exp VName ) )

Our infer catamorphism now returns an additional Reader (Pred -> Exp VName) (Exp VName) , which is to say an expression that has access to which expressions it should substitute for each of its Pred s. We will use this mapping to assign dictionaries to Pred s, allowing us to fill in the dictionary terms once we’ve figured them out.

We’re in the home stretch; now all we need to do is to have discharge build that map from Pred s into their dictionaries and we’re good to go.

getDictTerm :: Pred -> Exp VName getDictTypeForPred :: Pred -> Type -- DSL-level function application (:@) :: Exp VName -> Exp VName -> Exp VName discharge :: ClassEnv -> Pred -> TI ( Subst Pred ] , [ , Map Pred ( Exp VName ) Assump Type ] , [ Exp VName ] , [ ) = do discharge cenv p <- matchingInstances $ \(qs :=> t) -> do for (getInsts cenv)\(qst) <- ( fmap (qs, t, ) <$> match t p) <|> pure Nothing res(qs, t, )match t p) pure $ First res res case getFirst $ mconcat matchingInstances of getFirstmatchingInstances Just (qs, t, subst) -> (qs, t, subst) -- discharge all constraints on this instance (subst', qs', mapPreds, assumps, subDicts) <- fmap mconcat . traverse (discharge cenv) (discharge cenv) $ sub subst qs sub subst qs let dictTerm = getDictTerm t dictTermgetDictTerm t = foldl ( :@ ) dictTerm subDicts myDict) dictTerm subDicts pure ( subst' ( subst' , qs' <> M.singleton p myDict , mapPredsM.singleton p myDict , assumps -- this is just in a list so we can use 'mconcat' to -- collapse our traversal , [myDict] ) Nothing -> -- unable to discharge, so assume the existence of a new -- variable with the correct type <- newVName "d" paramnewVName pure ( mempty , [p] , M.singleton p param MkAssump param $ getDictTypeForPred p] , [paramgetDictTypeForPred p] , [param] )

The logic of discharge is largely the same, except we have a little more logic being driven by its new type. We now, in addition to our previous substitution and new predicates, also return a map expanding dictionaries, a list of Assump s (more on this in a second), and the resulting dictionary witnessing this discharged Pred .

If we were successful in finding a matching instance, we discharge each of its constraints, and fold the resulting dictionaries into ours. The more interesting logic is what happens if we are unable to discharge a constraint. In that case, we create a new variable of the necessary type, give that as our resulting dictionary, and emit it as an Assump . Assump s are used to denote the creation of a new variable in scope (they are also used for binding pattern matches).

The result of our new discharge function is that we have a map from every Pred we saw to the resulting dictionary for that instance, along with a list of generated variables. We can build our final expression tree via running the Reader (Pred -> Exp VName) by looking up the Pred s in our dictionary map. Finally, for every assumption we were left with, we fold our resulting term in a lambda which binds that assumption.

Very cool! If you’re interested in more of the nitty-gritty details behind compiling Haskell98, feel free to SMASH THAT STAR BUTTON on Github.