```
{-# LANGUAGE CPP #-}
#ifdef __HASTE__
import Haste.DOM
import Haste.Events
#else
import System.Console.Readline
#endif
import Control.Monad
import Data.Char
import Data.List
import Data.Tuple
import Text.ParserCombinators.Parsec
data Type = TV String | Forall String Type | Type :-> Type
data Term = Var String | App Term Term | Lam (String, Type) Term
| Let String Term Term
| TLam String Term | TApp Term Type
instance Show Type where
show (TV s) = s
show (Forall s t) = '\8704':(s ++ "." ++ show t)
show (t :-> u) = showL t ++ " -> " ++ showR u where
showL (Forall _ _) = "(" ++ show t ++ ")"
showL (_ :-> _) = "(" ++ show t ++ ")"
showL _ = show t
showR (Forall _ _) = "(" ++ show u ++ ")"
showR _ = show u
instance Show Term where
show (Lam (x, t) y) = "\0955" ++ x ++ showT t ++ showB y where
showB (Lam (x, t) y) = " " ++ x ++ showT t ++ showB y
showB expr = '.':show expr
showT (TV "_") = ""
showT t = ':':show t
show (TLam s t) = "\0955" ++ s ++ showB t where
showB (TLam s t) = " " ++ s ++ showB t
showB expr = '.':show expr
show (Var s) = s
show (App x y) = showL x ++ showR y where
showL (Lam _ _) = "(" ++ show x ++ ")"
showL _ = show x
showR (Var s) = ' ':s
showR _ = "(" ++ show y ++ ")"
show (TApp x y) = showL x ++ "[" ++ show y ++ "]" where
showL (Lam _ _) = "(" ++ show x ++ ")"
showL _ = show x
show (Let x y z) =
"let " ++ x ++ " = " ++ show y ++ " in " ++ show z
```

# System F

Simply typed lambda calculus is restrictive. The let-polymorphism of Hindley-Milner gives us more breathing room, but can we do better?

System F frees the type system further by introducing parts of lambda calculus
at the type level. We have *type abstraction* terms and *type application*
terms, which define and apply functions that take types as arguments and return
terms. At the same time, System F remains normalizing.

System F is rich enough that the self-application \x.x x is typable.

If we focus on the types, System F is a gentle generalization of Hindley-Milner. In the latter, any universal quantifiers (∀) must appear at the beginning of type, meaning they are scoped over the entire type. In System F, they can be arbitrary scoped. For example. (∀X.X->X)->(∀X.X->X) is a System F type, but not a Hindley-Milner type, while ∀X.(X->X)->X->X is a type in both systems.

This seemingly small change has deep consequences. Recall Hindley-Milner allows:

let id = \x.x in id succ(id 0)

but rejects:

(\f.f succ(f 0)) (\x.x)

This is because algorithm W assigns x a type variable, say X, then finds the conflicting constraints X = Nat and X = Nat -> Nat. A locally scoped generalized variable fixes this by causing fresh type variables to be generated for each use, so the resulting constraints, say, X = Nat and Y = Nat -> Nat, live and let live. Let-free let-polymorphism!

It’s easy to demonstrate this in Haskell. The following fails:

(\f->f succ(f 0))(\x->x)

We can make it run by enabling an extension, and annotating the identity function appropriately:

:set -XRankNTypes ((\f->f succ(f 0)) :: ((forall x.x->x)->Int))(\x->x)

We write type abstractions as lambda abstractions without a type annotation.
The simplest example is the *polymorphic identity function*:

id=\X.\x:X.x

For clarity, we capitalize the first letter of type variables in our examples.

The above represents a function that takes a type, and then returns an identity function for that type.

We write type applications like term applications except we surround the argument with square brackets. For example:

id [Nat] 42

evaluates to 42.

## Type spam

Our new features have a heavy price tag. In System F, type reconstruction is undecidable. We must add explicit types to every binding in every lambda abstraction. Moreover, applying type abstractions require us to state even more types.

Haskell magically fills in missing types if enough hints are given, which is why our example above got away with relatively little annotation. Our implementation of System F lacks this ability, so we must always supply types. (This is why we used Haskell above instead of our System F interpreter!)

Because types must frequently be specified, few practical languages are built on System F. (Perhaps it’s also because System F is a relatively recent discovery?) The Hindley-Milner system is often preferable, due to type inference.

However, behind the scenes, modern Haskell is an extension of System F. Certain language features require type annotation, and they generate unseen intermediate code full of type abstractions and type applications.

Type operators make types less eye-watering.

## Definitions

We replace the GV constructor representing Hindley-Milner generalized variables with its scoped version Forall.

We add type applications and type abstractions to the Term data type. A type application is like a term application, except it expects a type as input (and during printing should enclose it within square brackets). A type abstraction is like a term abstraction, except its variable, being a type variable, has no type annotation.

Universal types complicate type comparison, because bound type variables
may be renamed arbitrarily. That is, types are unique up to *alpha-conversion*.

```
instance Eq Type where
t1 == t2 = f [] t1 t2 where
f alpha (TV s) (TV t)
| Just t' <- lookup s alpha = t' == t
| Just _ <- lookup t (swap <$> alpha) = False
| otherwise = s == t
f alpha (Forall s x) (Forall t y)
| s == t = f alpha x y
| otherwise = f ((s, t):alpha) x y
f alpha (a :-> b) (c :-> d) = f alpha a c && f alpha b d
f alpha _ _ = False
```

## Parsing

Parsing is trickier because elements of lambda calculus have invaded the type level. For each abstraction, we look for a type binding to determine if it’s at the term or type level. For applications, we look for square brackets to decide.

We follow Haskell and accept forall as well as the harder-to-type ∀ symbol. Conventionally, the arrow type constructor -> has higher precedence.

We accept inputs that omit all but the last period and all but the first quantifier in a sequence of universal quantified type variables, an abbreviation similar to the one we’ve been using in sequences of abstractions.

```
data FLine = Empty | TopLet String Term | Run Term deriving Show
line :: Parser FLine
line = (((eof >>) . pure) =<<) . (ws >>) $ option Empty $
(try $ TopLet <$> v <*> (str "=" >> term)) <|> (Run <$> term) where
term = letx <|> lam <|> app
letx = Let <$> (str "let" >> v) <*> (str "=" >> term)
<*> (str "in" >> term)
lam = flip (foldr pickLam) <$> between lam0 lam1 (many1 vt) <*> term where
lam0 = str "\\" <|> str "\0955"
lam1 = str "."
vt = (,) <$> v <*> optionMaybe (str ":" >> typ)
pickLam (s, Nothing) = TLam s
pickLam (s, Just t) = Lam (s, t)
typ = forallt <|> fun
forallt = flip (foldr Forall) <$> between fa0 fa1 (many1 v) <*> fun where
fa0 = str "forall" <|> str "\8704"
fa1 = str "."
fun = ((TV <$> v)
<|> between (str "(") (str ")") typ) `chainr1` (str "->" >> pure (:->))
app = termArg >>= moreArg
termArg = (Var <$> v) <|> between (str "(") (str ")") term
moreArg t = option t $ ((App t <$> termArg)
<|> (TApp t <$> between (str "[") (str "]") typ)) >>= moreArg
v = try $ do
s <- many1 alphaNum
when (s `elem` words "let in forall") $ fail "unexpected keyword"
ws
pure s
str = try . (>> ws) . string
ws = spaces >> optional (try $ string "--" >> many anyChar)
```

## Typing

We’ve abandoned type inference, which simplifies typing. Nonetheless, we must handle the new terms. Type abstractions are trivial. As for type applications, once again, a routine used at the term level must now be written at the type level: we must rename type variables when they conflict with free type variables.

The shallow feature will be explained later.

```
typeOf :: [(String, Type)] -> Term -> Either String Type
typeOf gamma t = case t of
App (Var "shallow") y -> pure $ fst $ shallow gamma y
Var s | Just t <- lookup s gamma -> pure t
| otherwise -> Left $ "undefined " ++ s
App x y -> do
tx <- rec x
ty <- rec y
case tx of
ty' :-> tz | ty == ty' -> pure tz
_ -> Left $ show tx ++ " apply " ++ show ty
Lam (x, t) y -> do
u <- typeOf ((x, t):gamma) y
pure $ t :-> u
TLam s t -> Forall s <$> typeOf gamma t
TApp x y -> do
tx <- typeOf gamma x
case tx of
Forall s t -> pure $ beta t where
beta (TV v) | s == v = y
| otherwise = TV v
beta (Forall u v)
| s == u = Forall u v
| u `elem` fvs = let u1 = newName u fvs in
Forall u1 $ beta $ tRename u u1 v
| otherwise = Forall u $ beta v
beta (m :-> n) = beta m :-> beta n
fvs = tfv [] y
_ -> Left $ "TApp " ++ show tx
Let s t u -> do
tt <- typeOf gamma t
typeOf ((s, tt):gamma) u
where rec = typeOf gamma
tfv vs (TV s) | s `elem` vs = []
| otherwise = [s]
tfv vs (x :-> y) = tfv vs x `union` tfv vs y
tfv vs (Forall s t) = tfv (s:vs) t
tRename x x1 ty = case ty of
TV s | s == x -> TV x1
| otherwise -> ty
Forall s t
| s == x -> ty
| otherwise -> Forall s (rec t)
a :-> b -> rec a :-> rec b
where rec = tRename x x1
```

## Evaluation

Evaluation remains about the same. We erase types as we go.

As usual, the function eval takes us to weak head normal form:

```
eval env (App (Var "shallow") t) = snd $ shallow (fst env) t
eval env (Let x y z) = eval env $ beta (x, y) z
eval env (App m a) = let m' = eval env m in case m' of
Lam (v, _) f -> eval env $ beta (v, a) f
_ -> App m' a
eval env (TApp m _) = eval env m
eval env (TLam _ t) = eval env t
eval env term@(Var v) | Just x <- lookup v (snd env) = case x of
Var v' | v == v' -> x
_ -> eval env x
eval _ term = term
beta (v, a) f = case f of
Var s | s == v -> a
| otherwise -> Var s
Lam (s, _) m
| s == v -> Lam (s, TV "_") m
| s `elem` fvs -> let s1 = newName s fvs in
Lam (s1, TV "_") $ rec $ rename s s1 m
| otherwise -> Lam (s, TV "_") (rec m)
App m n -> App (rec m) (rec n)
TLam s t -> TLam s (rec t)
TApp t ty -> TApp (rec t) ty
Let x y z -> Let x (rec y) (rec z)
where
fvs = fv [] a
rec = beta (v, a)
fv vs (Var s) | s `elem` vs = []
| otherwise = [s]
fv vs (Lam (s, _) f) = fv (s:vs) f
fv vs (App x y) = fv vs x `union` fv vs y
fv vs (Let _ x y) = fv vs x `union` fv vs y
fv vs (TLam _ t) = fv vs t
fv vs (TApp x _) = fv vs x
newName x ys = head $ filter (`notElem` ys) $ (s ++) . show <$> [1..] where
s = dropWhileEnd isDigit x
rename x x1 term = case term of
Var s | s == x -> Var x1
| otherwise -> term
Lam (s, t) b
| s == x -> term
| otherwise -> Lam (s, t) (rec b)
App a b -> App (rec a) (rec b)
Let a b c -> Let a (rec b) (rec c)
TLam s t -> TLam s (rec t)
TApp a b -> TApp (rec a) b
where rec = rename x x1
```

The function norm recurses to find the normal form:

```
norm env@(gamma, lets) term = case eval env term of
Var v -> Var v
-- Record abstraction variable to avoid clashing with let definitions.
Lam (v, _) m -> Lam (v, TV "_") (norm (gamma, (v, Var v):lets) m)
App m n -> App (rec m) (rec n)
Let x y z -> Let x (rec y) (rec z)
TApp m _ -> rec m
TLam _ t -> rec t
where rec = norm env
```

## User Interface

The less said the better.

```
#ifdef __HASTE__
main = withElems ["input", "output", "evalB", "resetB", "resetP",
"pairB", "pairP", "surB", "surP"] $
\[iEl, oEl, evalB, resetB, resetP, pairB, pairP, surB, surP] -> do
let
reset = getProp resetP "value" >>= setProp iEl "value" >> setProp oEl "value" ""
run (out, env) (Left err) =
(out ++ "parse error: " ++ show err ++ "\n", env)
run (out, env@(gamma, lets)) (Right m) = case m of
Empty -> (out, env)
Run term -> case typeOf gamma term of
Left msg -> (out ++ "type error: " ++ msg ++ "\n", env)
Right t -> (out ++ show (norm env term) ++ "\n", env)
TopLet s term -> case typeOf gamma term of
Left msg -> (out ++ "type error: " ++ msg ++ "\n", env)
Right t -> (out ++ "[" ++ s ++ ":" ++ show t ++ "]\n",
((s, t):gamma, (s, term):lets))
reset
resetB `onEvent` Click $ const reset
pairB `onEvent` Click $ const $
getProp pairP "value" >>= setProp iEl "value" >> setProp oEl "value" ""
surB `onEvent` Click $ const $
getProp surP "value" >>= setProp iEl "value" >> setProp oEl "value" ""
evalB `onEvent` Click $ const $ do
es <- map (parse line "") . lines <$> getProp iEl "value"
setProp oEl "value" $ fst $ foldl' run ("", ([], [])) es
#else
repl env@(gamma, lets) = do
let redo = repl env
ms <- readline "> "
case ms of
Nothing -> putStrLn ""
Just s -> do
addHistory s
case parse line "" s of
Left err -> do
putStrLn $ "parse error: " ++ show err
redo
Right Empty -> redo
Right (Run term) ->
case typeOf gamma term of
Left msg -> putStrLn ("type error: " ++ msg) >> redo
Right t -> do
putStrLn $ "[type = " ++ show t ++ "]"
print $ norm env term
redo
Right (TopLet s term) -> case typeOf gamma term of
Left msg -> putStrLn ("type error: " ++ msg) >> redo
Right t -> do
putStrLn $ "[type = " ++ show t ++ "]"
repl ((s, t):gamma, (s, term):lets)
main = repl ([], [])
#endif
```

## Booleans, Integers, Pairs, Lists

Hindley-Milner supports Church-encodings of booleans, integers, pairs, and lists. System F is more general, so of course supports them too. However, we must be explicit about types. With Hindley-Milner, globally scoped universal quantifiers for all type variables are implied. With System F, our terms must start with type abstractions or a term abstraction annotated with a universal type:

true = \X x:X y:X.x false = \X x:X y:X.y not = \b:forall X.X->X->X X t:X f:X.b [X] f t 0 = \X s:X->X z:X.z succ = \n:(forall X.(X->X)->X->X) X s:X->X z:X.s(n[X] s z) pair = \X Y x:X y:Y Z f:X->Y->Z.f x y fst = \X Y p:forall Z.(X->Y->Z)->Z.p [X] (\x:X y:Y.x) snd = \X Y p:forall Z.(X->Y->Z)->Z.p [Y] (\x:X y:Y.y) nil = \X.(\R.\c:X->R->R.\n:R.n) cons = \X h:X t:forall R.(X->R->R)->R->R.(\R c:X->R->R n:R.c h (t [R] c n))

## Apply yourself!

In our previous systems, the term \x.x x has been untypable. No longer! Self-application can be expressed in System F:

\x:forall X.X->X.x[forall X.X->X] x

Of course, self-application of self-application ((\x.x x)(\x.x x)) remains untypable, because it has no normal form.

## Information hiding

It turns out universal types can represent *existential types*.
These types can enforce information hiding. For example, we can give
the programmer access to an API but forbid access to the implementation
details.

Knowledge is power. Languages with simpler type systems still benefit if their designers know System F. For example, Haskell 98 is strictly Hindley-Milner, but researchers exploited existential types to design and prove theorems about a language extension enabling the ST monad.

## System F in System F

The polymorphic identity function is typable in System F, hence we can
construct the self-interpreter described
In *Breaking Through the
Normalization Barrier: A Self-Interpreter for F-omega* by Matt Brown and Jens
Palsberg.

E=\T q:(forall X.X->X)->T.q(\X x:X.x)

As we demonstrated for the shallow encoding of untyped lambda calculus terms, the trick is to block every application (of types or terms) by adding one more layer of abstraction. The evaluation proceeds once we instantiate the abstraction with the polymorphic identity function.

Since we must specify types in System F, computing even a shallow encoding involves type checking.

```
shallow gamma term = (Forall "_0" (TV "_0" :-> TV "_0") :-> t,
Lam ("id", Forall "X" (TV "X" :-> TV "X")) q) where
(t, q) = f gamma term where
f gamma term = case term of
Var s -> (ty, Var s) where Just ty = lookup s gamma
App m n -> (tn, App (App (TApp (Var "id") tm) qm) qn) where
(tm, qm) = f gamma m
(tn, qn) = f gamma n
Lam (s, ty) t -> (ty :-> tt, Lam (s, ty) qt) where
(tt, qt) = f ((s, ty):gamma) t
TLam s t -> (Forall s tt, TLam s qt) where
(tt, qt) = f gamma t
TApp x ty -> (beta t, TApp (App (TApp (Var "id") tx) q) ty) where
(tx@(Forall s t), q) = f gamma x
beta (TV v) | s == v = ty
| otherwise = TV v
beta (Forall u v)
| s == u = Forall u v
| u `elem` fvs = let u1 = newName u fvs in
Forall u1 $ beta $ tRename u u1 v
| otherwise = Forall u $ beta v
beta (m :-> n) = beta m :-> beta n
fvs = tfv [] ty
Let s t u -> f ((s, fst $ f gamma t):gamma) $ beta (s, t) u
```

While this shallow construction is uninteresting, the existence of a self-interpreter for a strongly normalizing language is significant, as some books and papers claim this is impossible.

*blynn@cs.stanford.edu*💡