X-Git-Url: http://lambda.jimpryor.net/git/gitweb.cgi?p=lambda.git;a=blobdiff_plain;f=topics%2F_week5_system_F.mdwn;h=242ada430c4430e6076b668dbb6c4d819a6ad6a8;hp=a80cc58e340154930fc4114cb61e1eb62d3028ee;hb=be11c9fb4e8f436be20b5ccd5fb5a03794815440;hpb=5cab83962241676d710c788561ac107a3563a3e8
diff --git a/topics/_week5_system_F.mdwn b/topics/_week5_system_F.mdwn
index a80cc58e..242ada43 100644
--- a/topics/_week5_system_F.mdwn
+++ b/topics/_week5_system_F.mdwn
@@ -1,23 +1,13 @@
[[!toc levels=2]]
-# System F and recursive types
+# System F: the polymorphic lambda calculus
-In the simply-typed lambda calculus, we write types like σ
--> τ. This looks like logical implication. We'll take
-that resemblance seriously when we discuss the Curry-Howard
-correspondence. In the meantime, note that types respect modus
-ponens:
-
-
-Expression Type Implication
------------------------------------
-fn α -> β α ⊃ β
-arg α α
------- ------ --------
-(fn arg) β β
-
-
-The implication in the right-hand column is modus ponens, of course.
+The simply-typed lambda calculus is beautifully simple, but it can't
+even express the predecessor function, let alone full recursion. And
+we'll see shortly that there is good reason to be unsatisfied with the
+simply-typed lambda calculus as a way of expressing natural language
+meaning. So we will need to get more sophisticated about types. The
+next step in that journey will be to consider System F.
System F was discovered by Girard (the same guy who invented Linear
Logic), but it was independently proposed around the same time by
@@ -34,35 +24,34 @@ notational convention (which will last throughout the rest of the
course) that "x:α" represents an expression `x`
whose type is α.
-Then System F can be specified as follows (choosing notation that will
-match up with usage in O'Caml, whose type system is based on System F):
+Then System F can be specified as follows:
System F:
---------
- types Ï ::= c | α | Ï1 -> Ï2 | â'a. Ï
- expressions e ::= x | λx:Ï. e | e1 e2 | Îα. e | e [Ï]
+ types Ï ::= c | α | Ï1 -> Ï2 | âα.Ï
+ expressions e ::= x | λx:Ï.e | e1 e2 | Îα.e | e [Ï]
In the definition of the types, "`c`" is a type constant. Type
constants play the role in System F that base types play in the
simply-typed lambda calculus. So in a lingusitics context, type
-constants might include `e` and `t`. "α" is a type variable. The
-tick mark just indicates that the variable ranges over types rather
-than over values; in various discussion below and later, type variable
-can be distinguished by using letters from the greek alphabet
-(α, β, etc.), or by using capital roman letters (X, Y,
-etc.). "`Ï1 -> Ï2`" is the type of a function from expressions of
-type `Ï1` to expressions of type `Ï2`. And "`âα. Ï`" is called a
-universal type, since it universally quantifies over the type variable
-`'a`. You can expect that in `âα. Ï`, the type `Ï` will usually
-have at least one free occurrence of `α` somewhere inside of it.
+constants might include `e` and `t`. "α" is a type variable. In
+various discussions, type variables are distinguished by using letters
+from the greek alphabet (α, β, etc.), as we do here, or by
+using capital roman letters (X, Y, etc.), or by adding a tick mark
+(`'a`, `'b`, etc.), as in OCaml. "`Ï1 -> Ï2`" is the type of a
+function from expressions of type `Ï1` to expressions of type `Ï2`.
+And "`âα.Ï`" is called a universal type, since it universally
+quantifies over the type variable `α`. You can expect that in
+`âα.Ï`, the type `Ï` will usually have at least one free occurrence of
+`α` somewhere inside of it.
In the definition of the expressions, we have variables "`x`" as usual.
-Abstracts "`λx:Ï. e`" are similar to abstracts in the simply-typed lambda
+Abstracts "`λx:Ï.e`" are similar to abstracts in the simply-typed lambda
calculus, except that they have their shrug variable annotated with a
type. Applications "`e1 e2`" are just like in the simply-typed lambda calculus.
In addition to variables, abstracts, and applications, we have two
-additional ways of forming expressions: "`Îα. e`" is called a *type
+additional ways of forming expressions: "`Îα.e`" is called a *type
abstraction*, and "`e [Ï]`" is called a *type application*. The idea
is that Λ is a capital λ: just
like the lower-case λ, Λ binds
@@ -70,91 +59,87 @@ variables in its body, except that unlike λ,
Λ binds type variables instead of expression
variables. So in the expression
-Λ Î± (λ x:α . x)
+Λ Î± (λ x:α. x)
the Λ binds the type variable `α` that occurs in
-the λ abstract. Of course, as long as type
-variables are carefully distinguished from expression variables (by
-tick marks, Grecification, or capitalization), there is no need to
-distinguish expression abstraction from type abstraction by also
-changing the shape of the lambda.
+the λ abstract.
-The expression immediately below is a polymorphic version of the
-identity function. It defines one general identity function that can
-be adapted for use with expressions of any type. In order to get it
-ready to apply this identity function to, say, a variable of type
-boolean, just do this:
+This expression is a polymorphic version of the identity function. It
+defines one general identity function that can be adapted for use with
+expressions of any type. In order to get it ready to apply this
+identity function to, say, a variable of type boolean, just do this:
-(Λ Î± (λ x:α . x)) [t]
+(Λ Î± (λ x:α. x)) [t]
This type application (where `t` is a type constant for Boolean truth
values) specifies the value of the type variable `α`. Not
-surprisingly, the type of this type application is a function from
-Booleans to Booleans:
+surprisingly, the type of the expression that results from this type
+application is a function from Booleans to Booleans:
-((Λ Î± (λ x:α . x)) [t]): (b -> b)
+((Λα (λ x:α . x)) [t]): (b->b)
Likewise, if we had instantiated the type variable as an entity (base
type `e`), the resulting identity function would have been a function
of type `e -> e`:
-((Λ Î± (λ x:α . x)) [e]): (e -> e)
+((Λα (λ x:α. x)) [e]): (e->e)
Clearly, for any choice of a type `α`, the identity function can be
instantiated as a function from expresions of type `α` to expressions
of type `α`. In general, then, the type of the uninstantiated
(polymorphic) identity function is
-(Λ Î± (λ x:α . x)): (∀ α . α -> α)
+(Λα (λx:α . x)): (∀α. α->α)
Pred in System F
----------------
We saw that the predecessor function couldn't be expressed in the
simply-typed lambda calculus. It *can* be expressed in System F,
-however. Here is one way, coded in
-[[Benjamin Pierce's type-checker and evaluator for
-System F|http://www.cis.upenn.edu/~bcpierce/tapl/index.html]] (the
-relevant evaluator is called "fullpoly"):
-
- N = All X . (X->X)->X->X;
- Pair = (N -> N -> N) -> N;
- let zero = lambda X . lambda s:X->X . lambda z:X. z in
- let fst = lambda x:N . lambda y:N . x in
- let snd = lambda x:N . lambda y:N . y in
- let pair = lambda x:N . lambda y:N . lambda z:N->N->N . z x y in
- let suc = lambda n:N . lambda X . lambda s:X->X . lambda z:X . s (n [X] s z) in
- let shift = lambda p:Pair . pair (suc (p fst)) (p fst) in
- let pre = lambda n:N . n [Pair] shift (pair zero zero) snd in
+however. Here is one way:
+
+ let N = âα.(α->α)->α->α in
+ let Pair = (N->N->N)->N in
+
+ let zero = Îα. λs:α->α. λz:α. z in
+ let fst = λx:N. λy:N. x in
+ let snd = λx:N. λy:N. y in
+ let pair = λx:N. λy:N. λz:N->N->N. z x y in
+ let succ = λn:N. Îα. λs:α->α. λz:α. s (n [α] s z) in
+ let shift = λp:Pair. pair (succ (p fst)) (p fst) in
+ let pred = λn:N. n [Pair] shift (pair zero zero) snd in
pre (suc (suc (suc zero)));
-We've truncated the names of "suc(c)" and "pre(d)", since those are
-reserved words in Pierce's system. Note that in this code, there is
-no typographic distinction between ordinary lambda and type-level
-lambda, though the difference is encoded in whether the variables are
-lower case (for ordinary lambda) or upper case (for type-level
-lambda).
+[If you want to run this code in
+[[Benjamin Pierce's type-checker and evaluator for
+System F|http://www.cis.upenn.edu/~bcpierce/tapl/index.html]], the
+relevant evaluator is called "fullpoly", and you'll need to
+truncate the names of "suc(c)" and "pre(d)", since those are
+reserved words in Pierce's system.]
+
+Exercise: convince yourself that `zero` has type `N`.
The key to the extra expressive power provided by System F is evident
-in the typing imposed by the definition of `pre`. The variable `n` is
-typed as a Church number, i.e., as `All X . (X->X)->X->X`. The type
-application `n [Pair]` instantiates `n` in a way that allows it to
-manipulate ordered pairs: `n [Pair]: (Pair->Pair)->Pair->Pair`. In
-other words, the instantiation turns a Church number into a
+in the typing imposed by the definition of `pred`. The variable `n`
+is typed as a Church number, i.e., as `N ≡ âα.(α->α)->α->α`.
+The type application `n [Pair]` instantiates `n` in a way that allows
+it to manipulate ordered pairs: `n [Pair]: (Pair->Pair)->Pair->Pair`.
+In other words, the instantiation turns a Church number into a certain
pair-manipulating function, which is the heart of the strategy for
-this version of predecessor.
-
-Could we try to build a system for doing Church arithmetic in which
-the type for numbers always manipulated ordered pairs? The problem is
-that the ordered pairs we need here are pairs of numbers. If we tried
-to replace the type for Church numbers with a concrete (simple) type,
-we would have to replace each `X` with the type for Pairs, `(N -> N ->
-N) -> N`. But then we'd have to replace each of these `N`'s with the
-type for Church numbers, `(X -> X) -> X -> X`. And then we'd have to
-replace each of these `X`'s with... ad infinitum. If we had to choose
-a concrete type built entirely from explicit base types, we'd be
-unable to proceed.
+this version of computing the predecessor function.
+
+Could we try to accommodate the needs of the predecessor function by
+building a system for doing Church arithmetic in which the type for
+numbers always manipulated ordered pairs? The problem is that the
+ordered pairs we need here are pairs of numbers. If we tried to
+replace the type for Church numbers with a concrete (simple) type, we
+would have to replace each `N` with the type for Pairs, `(N -> N -> N)
+-> N`. But then we'd have to replace each of these `N`'s with the
+type for Church numbers, which we're imagining is `(Pair -> Pair) ->
+Pair -> Pair`. And then we'd have to replace each of these `Pairs`'s
+with... ad infinitum. If we had to choose a concrete type built
+entirely from explicit base types, we'd be unable to proceed.
[See Benjamin C. Pierce. 2002. *Types and Programming Languages*, MIT
Press, chapter 23.]
@@ -165,19 +150,17 @@ Typing ω
In fact, unlike in the simply-typed lambda calculus,
it is even possible to give a type for ω in System F.
-ω = lambda x:(All X. X->X) . x [All X . X->X] x
+ω = λx:(âα.α->α). x [âα.α->α] x
In order to see how this works, we'll apply ω to the identity
function.
-ω id ==
-
- (lambda x:(All X . X->X) . x [All X . X->X] x) (lambda X . lambda x:X . x)
+ω id ≡ (λx:(âα.α->α). x [âα.α->α] x) (Îα.λx:α.x)
-Since the type of the identity function is `(All X . X->X)`, it's the
+Since the type of the identity function is `âα.α->α`, it's the
right type to serve as the argument to ω. The definition of
ω instantiates the identity function by binding the type
-variable `X` to the universal type `All X . X->X`. Instantiating the
+variable `α` to the universal type `âα.α->α`. Instantiating the
identity function in this way results in an identity function whose
type is (in some sense, only accidentally) the same as the original
fully polymorphic identity function.
@@ -195,8 +178,8 @@ form in a finite number of steps.
Not only does a fixed-point combinator remain out of reach, we can't
even construct an infinite loop. This means that although we found a
type for ω, there is no general type for Ω ≡ ω
-ω. Furthermore, it turns out that no Turing complete system can
-be strongly normalizing, from which it follows that System F is not
+ω. In fact, it turns out that no Turing complete system can be
+strongly normalizing, from which it follows that System F is not
Turing complete.
@@ -209,10 +192,11 @@ The classic case study motivating polymorphism in natural language
comes from coordination. (The locus classicus is Partee and Rooth
1983.)
- Ann left and Bill left.
- Ann left and slept.
- Ann and Bill left.
- Ann read and reviewed the book.
+ Type of the argument of "and":
+ Ann left and Bill left. t
+ Ann left and slept. e->t
+ Ann and Bill left. (e->t)-t (i.e, generalize quantifiers)
+ Ann read and reviewed the book. e->e->t
In English (likewise, many other languages), *and* can coordinate
clauses, verb phrases, determiner phrases, transitive verbs, and many
@@ -220,7 +204,9 @@ other phrase types. In a garden-variety simply-typed grammar, each
kind of conjunct has a different semantic type, and so we would need
an independent rule for each one. Yet there is a strong intuition
that the contribution of *and* remains constant across all of these
-uses. Can we capture this using polymorphic types?
+uses.
+
+Can we capture this using polymorphic types?
Ann, Bill e
left, slept e -> t
@@ -228,41 +214,42 @@ uses. Can we capture this using polymorphic types?
With these basic types, we want to say something like this:
- and:t->t->t = lambda l:t . lambda r:t . l r false
- and = lambda 'a . lambda 'b .
- lambda l:'a->'b . lambda r:'a->'b .
- lambda x:'a . and:'b (l x) (r x)
-
-The idea is that the basic *and* conjoins expressions of type `t`, and
-when *and* conjoins functional types, it builds a function that
-distributes its argument across the two conjuncts and conjoins the two
-results. So `Ann left and slept` will evaluate to `(\x.and(left
-x)(slept x)) ann`. Following the terminology of Partee and Rooth, the
-strategy of defining the coordination of expressions with complex
-types in terms of the coordination of expressions with less complex
-types is known as Generalized Coordination.
-
-But the definitions just given are not well-formed expressions in
-System F. There are three problems. The first is that we have two
-definitions of the same word. The intention is for one of the
-definitions to be operative when the type of its arguments is type
-`t`, but we have no way of conditioning evaluation on the *type* of an
-argument. The second is that for the polymorphic definition, the term
-*and* occurs inside of the definition. System F does not have
-recursion.
-
-The third problem is more subtle. The defintion as given takes two
-types as parameters: the type of the first argument expected by each
-conjunct, and the type of the result of applying each conjunct to an
-argument of that type. We would like to instantiate the recursive use
-of *and* in the definition by using the result type. But fully
-instantiating the definition as given requires type application to a
-pair of types, not to just a single type. We want to somehow
-guarantee that 'b will always itself be a complex type.
-
-So conjunction and disjunction provide a compelling motivation for
-polymorphism in natural language, but we don't yet have the ability to
-build the polymorphism into a formal system.
+ and:t->t->t = λl:t. λr:t. l r false
+ gen_and = Îα.Îβ.λf:(β->t).λl:α->β.λr:α->β.λx:α. f (l x) (r x)
+
+The idea is that the basic *and* (the one defined in the first line)
+conjoins expressions of type `t`. But when *and* conjoins functional
+types (the definition in the second line), it builds a function that
+distributes its argument across the two conjuncts and then applies the
+appropriate lower-order instance of *and*.
+
+ and (Ann left) (Bill left)
+ gen_and [e] [t] and left slept
+ gen_and [e] [e->t] (gen_and [e] [t] and) read reviewed
+
+Following the terminology of Partee and Rooth, this strategy of
+defining the coordination of expressions with complex types in terms
+of the coordination of expressions with less complex types is known as
+Generalized Coordination, which is why we call the polymorphic part of
+the definition `gen_and`.
+
+In the first line, the basic *and* is ready to conjoin two truth
+values. In the second line, the polymorphic definition of `gen_and`
+makes explicit exactly how the meaning of *and* when it coordinates
+verb phrases depends on the meaning of the basic truth connective.
+Likewise, when *and* coordinates transitive verbs of type `e->e->t`,
+the generalized *and* depends on the `e->t` version constructed for
+dealing with coordinated verb phrases.
+
+On the one hand, this definition accurately expresses the way in which
+the meaning of the conjunction of more complex types relates to the
+meaning of the conjunction of simpler types. On the other hand, it's
+awkward to have to explicitly supply an expression each time that
+builds up the meaning of the *and* that coordinates the expressions of
+the simpler types. We'd like to have that automatically handled by
+the polymorphic definition; but that would require writing code that
+behaved differently depending on the types of its type arguments,
+which goes beyond the expressive power of System F.
And in fact, discussions of generalized coordination in the
linguistics literature are almost always left as a meta-level
@@ -271,6 +258,10 @@ Hendriks' 1992:74 dissertation, generalized coordination is
implemented as a method for generating a suitable set of translation
rules, which are in turn expressed in a simply-typed grammar.
+There is some work using System F to express generalizations about
+natural language: Ponvert, Elias. 2005. Polymorphism in English Logical
+Grammar. In *Lambda Calculus Type Theory and Natural Language*: 47--60.
+
Not incidentally, we're not aware of any programming language that
makes generalized coordination available, despite is naturalness and
ubiquity in natural language. That is, coordination in programming