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Haskell代写 | COMP3161/9164 20T3 Assignment 2

这个作业是用Haskell实现MinHS的类型推断过程
COMP3161/9164 20T3 Assignment 2
Type Inference for Polymorphic MinHs

1 Task 1
Task 1 is worth 100% of the marks of this assignment. You are to implement type inference for
MinHS with aggregate data structures. The following cases must be handled:
• the MinHS language of the first task of assignment 1 (without n-ary functions or letrecs, or
lists)
• product types: the 0-tuple and 2-tuples.
• sum types
1
• polymorphic functions
These cases are explained in detail below. The abstract syntax defining these syntactic entities is
in Syntax.hs. You should not need to modify the abstract syntax definition in any way.
Your implementation is to follow the definition of algebraic data types found in the lecture on
data structures, and the rules defined in the lectures on type inference. Additional material can be
found in the lecture notes on polymorphism, and the reference materials. The full set of rules is
outlined in this specification.
2 Bonus Tasks
These tasks are all optional, and are worth a total of an additional 10%.
2.1 Bonus Task 1: Simple Syntax Extensions
This bonus task is worth an additional 3%. In this task, you should implement type inference for
n-ary functions, and multiple bindings in the one let expression, with the same semantics as the
extension tasks for Assignment 1.
You will need to develop the requisite extensions to the type inference algorithm yourself, but
the extensions are very similar to the existing rules.
2.2 Bonus Task 2: Recursive Bindings
This bonus task is worth an additional 3%. In this task you must implement type inference for
letrec constructs.
Each binding will require a fresh unification variable, similar to multiple-binding let expressions,
however for letrec we do not wish types to be generalised.
2.3 Bonus Task 3: User-provided type signatures
This bonus task is worth an additional 4%. In this task you are to extend the type inference pass
to accept programs containing some type information. You need to combine this with the results
of your type inference pass to produce the final type for each declaration. That is, you need to be
able to infer correct types for programs like:
main = let f :: (Int -> Int)
= recfun g x = x;
in f 2;
You must ensure that the type signatures provided are not overly general. For example, the following
program should be a type error, as the user-provided signature is too general:
main :: (forall a. a) = 3;
3 Algebraic Data Types
This section covers the extensions to the language of the first assignment. In all other respects
(except lists) the language remains the same, so you can consult the reference material from the
first assignment for further details on the language.
2
3.1 Product Types
We only have 2-tuples in MinHS, and the Unit type, which could be viewed as a 0 type.
Types τ → τ 1 * τ 2
| 1
Expressions exp → (e1, e2)
| fst e | snd e
| ()
3.2 Sum Types
Types τ → τ + τ
Expressions exp → Inl e1
| Inr e2
| case e of
Inl x -> e1 ;
Inr y -> e2 ;
3.3 Polymorphism
The extensions to allow polymorphism are relatively minor. Two new type forms have been introduced: the TypeVar t form, and the Forall t e form.
Types τ → forall τ. ..τ..
For example, consider the following code fragment before and after type inference:
main =
let f = recfun g x = x;
in if f True
then f (Inl 1)
else f (Inr ());
main :: (Int + 1) =
let f :: forall a. (a -> a) = recfun g :: (a -> a) x = x;
in if f True
then f (Inl 1)
else f (Inr ());
3
4 Type Inference Rules
Sections coloured in blue can be viewed as inputs to the algorithm, while red text can be viewed
as outputs.
Constants and Variables
Γ ` n : Int
x : ∀a1 . . . ∀an.τ ∈ Γ
Γ ` x : τ [a1 := β1, . . . , an := βn]
βi fresh
(unquantify in TyInfer.hs provides an easy way to implement this)
Constructors and Primops
constructorType(C) = ∀a1 . . . ∀an.τ
Γ ` C : τ [a1 := β1, . . . , an := βn]
primopType(o) = ∀a1 . . . ∀an.τ
Γ ` o : τ [a1 := β1, . . . , an := βn]
βi fresh
(constructorType and primopType are defined in TyInfer.hs)
Application
TΓ ` e1 : τ1 T
0TΓ ` e2 : τ2 T
0
τ1
U∼ (τ2 → α)
UT0TΓ ` Apply e1 e2 : Uα
α fresh
If-Then-Else
TΓ ` e : τ τ
U∼ Bool T1UTΓ ` e1 : τ1 T2T1UTΓ ` e2 : τ2 T2τ1
U
0
∼ τ2
U0T2T1UTΓ ` If e e1 e2 : U0τ2
Case
TΓ ` e : τ
T1T(Γ ∪ {x : αl}) ` e1 : τl
T2T1T(Γ ∪ {y : αr}) ` e2 : τr
T2T1T(αl + αr)
U∼ T2T1τ UT2τl
U
0
∼ Uτr
U0UT2T1TΓ ` Case e x.e1 y.e2 : U0Uτr
αl
, αr fresh
Recursive Functions
T(Γ ∪ {x : α1} ∪ {f : α2}) ` e : τ T α2
U∼ T α1 → τ
UTΓ ` Recfun f.x.e : U(T α1 → τ )
α1, α2 fresh
Let Bindings
TΓ ` e1 : τ T
0
(TΓ ∪ {x : Generalise(TΓ, τ )}) ` e2 : τ
0
T0TΓ ` Let e1 x.e2 : τ
0
TΓ ` e : τ
Γ ` Main e : Generalise(Γ, τ )
where Generalise(Γ, τ ) = ∀ (TV (τ ) \ TV (Γ)) . τ
4.1 Implementing the algorithm
The type inference rules imply an algorithm, where the expression and the environment can be seen
as input, and the substitution and the type of the expression can be seen as output, as seen from
the color coding above.
Such an algorithm would probably have a type signature like:
4
inferExp :: Gamma -> Exp -> TC (Type, Subst)
However our goal isn’t just to infer the type of the expression, but to also elaborate the expression to
its explicitly-typed equivalent. In other words, we want to add typing annotations to our expressions,
meaning we must return a new expression as well as the type and substitution.
inferExp :: Gamma -> Exp -> TC (Exp, Type, Subst)
The cases for let and recfun (and letrec, in the bonus) must add a (correct) type signature to
the resultant expression, and all other cases must be sure to return a new expression consisting of
the updated subexpressions.
Because we add annotations to the expressions on the way, those type annotations wouldn’t
contain the information we get from typing successive expressions. Consider the following example:
{+ : Int → Int → Int, x : a} ` ( (let z = x in x) , x+1)
When typing let z = x in x, we only know that x is of type a, so we would add the type annotation let z :: a = x in x. Only when we type the second expression, x+1, we know that a has
to be Int (which will be reflected in T
0
). That’s why, when we’re done typing the top-level binding,
we have to traverse the whole binding again, applying the substitution to each type annotation
anywhere in the binding.
The function allTypes, defined in Syntax.hs, allows you to update each type annotation in an
expression, making it easy to perform this substitution.
5 Substitution
Substitutions are implemented as an abstract data type, defined in Subst.hs. Subst is an instance
of the Monoid type class, which is defined in the standard library as follows:
class Monoid a where
mappend :: a -> a -> a — also written as the infix operator <>
mempty :: a
For the Subst instance, mempty corresponds to the empty substitution, and mappend is substitution composition. That is, applying the substitution a <> b is the same as applying a and b
simultaneously. It should be reasonably clear that this instance obeys the monoid laws:
mempty <> x == x — left identity
x <> mempty == x — right identity
x <> (y <> z) == (x <> y) <> z — associativity
It is also commutative (x <> y == y <> x) assuming that the substitutions are disjoint (i.e that
dom(x) ∩ dom(y) = ∅). In the type inference algorithm, your substitutions are all applied in order
and thus should be disjoint, therefore this property should hold.
You can use this <> operator to combine multiple substitutions into your return substitution.
You can construct a singleton substitution, which replaces one variable, with the =: operator,
so the substitution (“a” =: TypeVar “b”) <> (“b” =: TypeVar “c”) is a substitution which
replaces a with c and b with c.
The Subst module also includes a variety of functions for running substitutions on types, quantified types, and environments.
5
6 Unification
The unification algorithm is the structural matching one originally due to Robinson and discussed
in lectures:
• input: two type terms t1 and t2, where forall quantified variables have been replaced by fresh,
unique variables
• output: the most general unifier of t1 and t2 (if it exists). The unifier is a data structure or
function that specifies the substitutions to take place to unify t1 and t2.
6.1 Unification Cases
For t1 and t2
1. both are type variables v1 and v2:
– if v1 = v2, return the empty substitution
– otherwise, return [v1 := v2]
2. both are primitive types
– if they are the same, return the empty substitution
– otherwise, there is no unifier
3. both are product types, with t1 = (t11 ∗ t12), t1 = (t21 ∗ t22)
– compute the unifier S of t11 and t21
– compute the unifier S
0 of S t12 and S t22
– return S <> S
0
4. function types and sum types (as for product types)
5. only one is a type variable v , the other an arbitrary type term t
– if v occurs in t, there is no unifier
– otherwise, return [v := t]
6. otherwise, there is no unifier
Functions in the Data.List library are useful for implementing the occurs check.
Once you generate a unifier (also a substitution), you then need to apply that unifier to your
types, to produce the unified type.
7 Errors and Fresh Names
Thus far, the following type signature would be sufficient for implementing our type inference
function:
inferExp :: Gamma -> Exp -> (Exp, Type, Subst)
Unification is a partial function, however, so we need well-founded way to handle the error cases,
rather than just bail out with error calls.
To achieve this, we’ll adjust the basic, pure signature for type inference to include the possibility
of a TypeError:
6
inferExp :: Gamma -> Exp -> Either TypeError (Exp, Type, Subst)
Even this, though, is not sufficient, as we cannot generate fresh, unique type variables for use as
unification variables:
fresh :: Type — it is impossible for fresh to return a different
fresh = ? — value each time!
To solve this problem, we could pass an (infinite) list of unique names around our program, and
fresh could simply take a name from the top of the list, and return a new list with the name
removed:
fresh :: [Id] -> ([Id],Type)
fresh (x:xs) = (xs,TypeVar x)
This is quite awkward though, as now we have to manually thread our list of identifiers throughout
the entire inference algorithm:
inferExp :: Gamma -> Exp -> [Id] -> Either TypeError ([Id], (Exp, Type, Subst))
To resolve this, we bundle both the [Id] state transformer and the Either TypeError x error
handling into one abstract type, called TC (defined in TCMonad.hs)
newtype TC a = TC ([Id] -> Either TypeError ([Id], a))
One can think of TC a abstractly as referring to a stateful action that will, if executed, return a
value of type a, or throw an exception.
As the name of the module implies, TC is a Monad, meaning that it exposes two functions
(return and >>=) as part of its interface.
return :: a -> TC a
return = …
(>>) :: TC a -> TC b -> TC b
a >> b = a >>= const b
(>>=) :: TC a -> (a -> TC b) -> TC b
a >>= b = …
The function return is, despite its name, just an ordinary function which lifts pure values into a TC
action that returns that value. The function (>>) (read then), is a kind of composition operator,
which produces a TC action which runs two TC actions in sequence, one after the other, returning
the value of the last one to be executed. Lastly, the function (>>=), more general than (>>), allows
the second executed action to be determined by the return value of the first.
The TCMonad.hs module also includes a few built-in actions:
typeError :: TypeError -> TC a — throw an error
fresh :: TC Type — return a fresh type variable
Haskell includes special syntactic sugar for monads, which allow for programming in a somewhat
imperative style. Called do notation, it is simple sugar for (>>) and (>>=).
do e — e
do e; v — e >> do v
do p <- e; v -- e >>= \p -> do v
do let x = y; v — let x = y in do v
This lets us write unification and type inference quite naturally. A simple example of the use of the
TC monad is already provided to you, with the unquantify function, which takes a type with some
number of quantifiers, and replaces all quantifiers with fresh variables (very useful for inference
cases for variables, constructors, and primops):
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unquantify :: QType -> TC Type
unquantify (Ty t) = return t
unquantify (Forall x t) = do x’ <- fresh unquantify (substQType (x =:x’) t) To run a TC action, in your top level infer function, the runTC function can be used: runTC :: TC a -> Either TypeError a
Please note: This function runs the TC action with the same source of fresh names each time!
Using it more than once in your program is not likely to give correct results.


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