(** * MoreStlc: More on the Simply Typed Lambda-Calculus *)
Require Import Maps.
Require Import Types.
Require Import Smallstep.
Require Import Stlc.
(* ################################################################# *)
(** * Simple Extensions to STLC *)
(** The simply typed lambda-calculus has enough structure to make its
theoretical properties interesting, but it is not much of a
programming language.
In this chapter, we begin to close the gap with real-world
languages by introducing a number of familiar features that have
straightforward treatments at the level of typing. *)
(* ================================================================= *)
(** ** Numbers *)
(** As we saw in exercise [stlc_arith] at the end of the [StlcProp]
chapter, adding types, constants, and primitive operations for
natural numbers is easy -- basically just a matter of combining
the [Types] and [Stlc] chapters. Adding more realistic
numeric types like machine integers and floats is also
straightforward, though of course the specifications of the
numeric primitives become more fiddly. *)
(* ================================================================= *)
(** ** Let Bindings *)
(** When writing a complex expression, it is useful to be able
to give names to some of its subexpressions to avoid repetition
and increase readability. Most languages provide one or more ways
of doing this. In OCaml (and Coq), for example, we can write [let
x=t1 in t2] to mean "reduce the expression [t1] to a value and
bind the name [x] to this value while reducing [t2]."
Our [let]-binder follows OCaml in choosing a standard
_call-by-value_ evaluation order, where the [let]-bound term must
be fully reduced before reduction of the [let]-body can begin.
The typing rule [T_Let] tells us that the type of a [let] can be
calculated by calculating the type of the [let]-bound term,
extending the context with a binding with this type, and in this
enriched context calculating the type of the body (which is then
the type of the whole [let] expression).
At this point in the book, it's probably easier simply to look at
the rules defining this new feature than to wade through a lot of
English text conveying the same information. Here they are: *)
(** Syntax:
t ::= Terms
| ... (other terms same as before)
| let x=t in t let-binding
*)
(**
Reduction:
t1 ==> t1'
---------------------------------- (ST_Let1)
let x=t1 in t2 ==> let x=t1' in t2
---------------------------- (ST_LetValue)
let x=v1 in t2 ==> [x:=v1]t2
Typing:
Gamma |- t1 : T1 Gamma, x:T1 |- t2 : T2
-------------------------------------------- (T_Let)
Gamma |- let x=t1 in t2 : T2
*)
(* ================================================================= *)
(** ** Pairs *)
(** Our functional programming examples in Coq have made
frequent use of _pairs_ of values. The type of such a pair is
called a _product type_.
The formalization of pairs is almost too simple to be worth
discussing. However, let's look briefly at the various parts of
the definition to emphasize the common pattern. *)
(** In Coq, the primitive way of extracting the components of a pair
is _pattern matching_. An alternative is to take [fst] and
[snd] -- the first- and second-projection operators -- as
primitives. Just for fun, let's do our pairs this way. For
example, here's how we'd write a function that takes a pair of
numbers and returns the pair of their sum and difference:
\x : Nat*Nat.
let sum = x.fst + x.snd in
let diff = x.fst - x.snd in
(sum,diff)
*)
(** Adding pairs to the simply typed lambda-calculus, then, involves
adding two new forms of term -- pairing, written [(t1,t2)], and
projection, written [t.fst] for the first projection from [t] and
[t.snd] for the second projection -- plus one new type constructor,
[T1*T2], called the _product_ of [T1] and [T2]. *)
(** Syntax:
t ::= Terms
| (t,t) pair
| t.fst first projection
| t.snd second projection
| ...
v ::= Values
| (v,v) pair value
| ...
T ::= Types
| T * T product type
| ...
*)
(** For reduction, we need several new rules specifying how pairs and
projection behave. *)
(**
t1 ==> t1'
-------------------- (ST_Pair1)
(t1,t2) ==> (t1',t2)
t2 ==> t2'
-------------------- (ST_Pair2)
(v1,t2) ==> (v1,t2')
t1 ==> t1'
------------------ (ST_Fst1)
t1.fst ==> t1'.fst
------------------ (ST_FstPair)
(v1,v2).fst ==> v1
t1 ==> t1'
------------------ (ST_Snd1)
t1.snd ==> t1'.snd
------------------ (ST_SndPair)
(v1,v2).snd ==> v2
*)
(** Rules [ST_FstPair] and [ST_SndPair] say that, when a fully
reduced pair meets a first or second projection, the result is
the appropriate component. The congruence rules [ST_Fst1] and
[ST_Snd1] allow reduction to proceed under projections, when the
term being projected from has not yet been fully reduced.
[ST_Pair1] and [ST_Pair2] reduce the parts of pairs: first the
left part, and then -- when a value appears on the left -- the right
part. The ordering arising from the use of the metavariables [v]
and [t] in these rules enforces a left-to-right evaluation
strategy for pairs. (Note the implicit convention that
metavariables like [v] and [v1] can only denote values.) We've
also added a clause to the definition of values, above, specifying
that [(v1,v2)] is a value. The fact that the components of a pair
value must themselves be values ensures that a pair passed as an
argument to a function will be fully reduced before the function
body starts executing. *)
(** The typing rules for pairs and projections are straightforward. *)
(**
Gamma |- t1 : T1 Gamma |- t2 : T2
--------------------------------------- (T_Pair)
Gamma |- (t1,t2) : T1*T2
Gamma |- t1 : T11*T12
--------------------- (T_Fst)
Gamma |- t1.fst : T11
Gamma |- t1 : T11*T12
--------------------- (T_Snd)
Gamma |- t1.snd : T12
*)
(** [T_Pair] says that [(t1,t2)] has type [T1*T2] if [t1] has
type [T1] and [t2] has type [T2]. Conversely, [T_Fst] and [T_Snd]
tell us that, if [t1] has a product type [T11*T12] (i.e., if it
will reduce to a pair), then the types of the projections from
this pair are [T11] and [T12]. *)
(* ================================================================= *)
(** ** Unit *)
(** Another handy base type, found especially in languages in
the ML family, is the singleton type [Unit]. *)
(** It has a single element -- the term constant [unit] (with a small
[u]) -- and a typing rule making [unit] an element of [Unit]. We
also add [unit] to the set of possible values -- indeed, [unit] is
the _only_ possible result of reducing an expression of type
[Unit]. *)
(** Syntax:
t ::= Terms
| unit unit value
| ...
v ::= Values
| unit unit
| ...
T ::= Types
| Unit Unit type
| ...
Typing:
-------------------- (T_Unit)
Gamma |- unit : Unit
*)
(** It may seem a little strange to bother defining a type that
has just one element -- after all, wouldn't every computation
living in such a type be trivial?
This is a fair question, and indeed in the STLC the [Unit] type is
not especially critical (though we'll see two uses for it below).
Where [Unit] really comes in handy is in richer languages with
_side effects_ -- e.g., assignment statements that mutate
variables or pointers, exceptions and other sorts of nonlocal
control structures, etc. In such languages, it is convenient to
have a type for the (trivial) result of an expression that is
evaluated only for its effect. *)
(* ================================================================= *)
(** ** Sums *)
(** Many programs need to deal with values that can take two distinct
forms. For example, we might identify employees in an accounting
application using _either_ their name _or_ their id number.
A search function might return _either_ a matching value _or_ an
error code.
These are specific examples of a binary _sum type_ (sometimes called
a _disjoint union_), which describes a set of values drawn from
one of two given types, e.g.:
Nat + Bool
*)
(** We create elements of these types by _tagging_ elements of
the component types. For example, if [n] is a [Nat] then [inl n]
is an element of [Nat+Bool]; similarly, if [b] is a [Bool] then
[inr b] is a [Nat+Bool]. The names of the tags [inl] and [inr]
arise from thinking of them as functions
inl : Nat -> Nat + Bool
inr : Bool -> Nat + Bool
that "inject" elements of [Nat] or [Bool] into the left and right
components of the sum type [Nat+Bool]. (But note that we don't
actually treat them as functions in the way we formalize them:
[inl] and [inr] are keywords, and [inl t] and [inr t] are primitive
syntactic forms, not function applications.) *)
(** In general, the elements of a type [T1 + T2] consist of the
elements of [T1] tagged with the token [inl], plus the elements of
[T2] tagged with [inr]. *)
(** One important usage of sums is signaling errors:
div : Nat -> Nat -> (Nat + Unit) =
div =
\x:Nat. \y:Nat.
if iszero y then
inr unit
else
inl ...
*)
(** The type [Nat + Unit] above is in fact isomorphic to [option
nat] in Coq -- i.e., it's easy to write functions that translate
back and forth. *)
(** To _use_ elements of sum types, we introduce a [case]
construct (a very simplified form of Coq's [match]) to destruct
them. For example, the following procedure converts a [Nat+Bool]
into a [Nat]: *)
(**
getNat =
\x:Nat+Bool.
case x of
inl n => n
| inr b => if b then 1 else 0
*)
(** More formally... *)
(** Syntax:
t ::= Terms
| inl T t tagging (left)
| inr T t tagging (right)
| case t of case
inl x => t
| inr x => t
| ...
v ::= Values
| inl T v tagged value (left)
| inr T v tagged value (right)
| ...
T ::= Types
| T + T sum type
| ...
*)
(** Reduction:
t1 ==> t1'
---------------------- (ST_Inl)
inl T t1 ==> inl T t1'
t1 ==> t1'
---------------------- (ST_Inr)
inr T t1 ==> inr T t1'
t0 ==> t0'
------------------------------------------- (ST_Case)
case t0 of inl x1 => t1 | inr x2 => t2 ==>
case t0' of inl x1 => t1 | inr x2 => t2
---------------------------------------------- (ST_CaseInl)
case (inl T v0) of inl x1 => t1 | inr x2 => t2
==> [x1:=v0]t1
---------------------------------------------- (ST_CaseInr)
case (inr T v0) of inl x1 => t1 | inr x2 => t2
==> [x2:=v0]t2
*)
(** Typing:
Gamma |- t1 : T1
---------------------------- (T_Inl)
Gamma |- inl T2 t1 : T1 + T2
Gamma |- t1 : T2
---------------------------- (T_Inr)
Gamma |- inr T1 t1 : T1 + T2
Gamma |- t0 : T1+T2
Gamma , x1:T1 |- t1 : T
Gamma , x2:T2 |- t2 : T
--------------------------------------------------- (T_Case)
Gamma |- case t0 of inl x1 => t1 | inr x2 => t2 : T
We use the type annotation in [inl] and [inr] to make the typing
relation simpler, similarly to what we did for functions. *)
(** Without this extra information, the typing rule [T_Inl], for
example, would have to say that, once we have shown that [t1] is
an element of type [T1], we can derive that [inl t1] is an element
of [T1 + T2] for _any_ type T2. For example, we could derive both
[inl 5 : Nat + Nat] and [inl 5 : Nat + Bool] (and infinitely many
other types). This peculiarity (technically, a failure of
uniqueness of types) would mean that we cannot build a
typechecking algorithm simply by "reading the rules from bottom to
top" as we could for all the other features seen so far.
There are various ways to deal with this difficulty. One simple
one -- which we've adopted here -- forces the programmer to
explicitly annotate the "other side" of a sum type when performing
an injection. This is a bit heavy for programmers (so real
languages adopt other solutions), but it is easy to understand and
formalize. *)
(* ================================================================= *)
(** ** Lists *)
(** The typing features we have seen can be classified into _base
types_ like [Bool], and _type constructors_ like [->] and [*] that
build new types from old ones. Another useful type constructor is
[List]. For every type [T], the type [List T] describes
finite-length lists whose elements are drawn from [T].
In principle, we could encode lists using pairs, sums and
_recursive_ types. But giving semantics to recursive types is
non-trivial. Instead, we'll just discuss the special case of lists
directly.
Below we give the syntax, semantics, and typing rules for lists.
Except for the fact that explicit type annotations are mandatory
on [nil] and cannot appear on [cons], these lists are essentially
identical to those we built in Coq. We use [lcase] to destruct
lists, to avoid dealing with questions like "what is the [head] of
the empty list?" *)
(** For example, here is a function that calculates the sum of
the first two elements of a list of numbers:
\x:List Nat.
lcase x of nil -> 0
| a::x' -> lcase x' of nil -> a
| b::x'' -> a+b
*)
(**
Syntax:
t ::= Terms
| nil T
| cons t t
| lcase t of nil -> t | x::x -> t
| ...
v ::= Values
| nil T nil value
| cons v v cons value
| ...
T ::= Types
| List T list of Ts
| ...
*)
(** Reduction:
t1 ==> t1'
-------------------------- (ST_Cons1)
cons t1 t2 ==> cons t1' t2
t2 ==> t2'
-------------------------- (ST_Cons2)
cons v1 t2 ==> cons v1 t2'
t1 ==> t1'
---------------------------------------- (ST_Lcase1)
(lcase t1 of nil -> t2 | xh::xt -> t3) ==>
(lcase t1' of nil -> t2 | xh::xt -> t3)
----------------------------------------- (ST_LcaseNil)
(lcase nil T of nil -> t2 | xh::xt -> t3)
==> t2
----------------------------------------------- (ST_LcaseCons)
(lcase (cons vh vt) of nil -> t2 | xh::xt -> t3)
==> [xh:=vh,xt:=vt]t3
*)
(** Typing:
----------------------- (T_Nil)
Gamma |- nil T : List T
Gamma |- t1 : T Gamma |- t2 : List T
----------------------------------------- (T_Cons)
Gamma |- cons t1 t2: List T
Gamma |- t1 : List T1
Gamma |- t2 : T
Gamma , h:T1, t:List T1 |- t3 : T
------------------------------------------------- (T_Lcase)
Gamma |- (lcase t1 of nil -> t2 | h::t -> t3) : T
*)
(* ================================================================= *)
(** ** General Recursion *)
(** Another facility found in most programming languages (including
Coq) is the ability to define recursive functions. For example,
we might like to be able to define the factorial function like
this:
fact = \x:Nat.
if x=0 then 1 else x * (fact (pred x)))
Note that the right-hand side of this binder mentions the variable
being bound -- something that is not allowed by our formalization of
[let] above.
Directly formalizing this "recursive definition" mechanism is possible,
but it requires a bit of extra effort: in particular, we'd have to
pass around an "environment" of recursive function definitions in
the definition of the [step] relation. *)
(** Here is another way of presenting recursive functions that is equally
powerful (though not quite as convenient for the programmer) and
more straightforward to formalize: instead of writing recursive
definitions, we define a _fixed-point operator_ called [fix]
that performs the "unfolding" of the recursive definition in the
right-hand side as needed, during reduction.
For example, instead of
fact = \x:Nat.
if x=0 then 1 else x * (fact (pred x)))
we will write:
fact =
fix
(\f:Nat->Nat.
\x:Nat.
if x=0 then 1 else x * (f (pred x)))
We can derive the latter from the former as follows:
- In the right-hand side of the definition of [fact], replace
recursive references to [fact] by a fresh variable [f].
- Add an abstraction binding [f] at the front, with an
appropriate type annotation. (Since we are using [f] in place
of [fact], which had type [Nat->Nat], we should require [f]
to have the same type.) The new abstraction has type
[(Nat->Nat) -> (Nat->Nat)].
- Apply [fix] to this abstraction. This application has
type [Nat->Nat].
- Use all of this as the right-hand side of an ordinary
[let]-binding for [fact].
*)
(** The intuition is that the higher-order function [f] passed
to [fix] is a _generator_ for the [fact] function: if [f] is
applied to a function that "approximates" the desired behavior of
[fact] up to some number [n] (that is, a function that returns
correct results on inputs less than or equal to [n] but we don't
care what it does on inputs greater than [n]), then [f] returns a
slightly better approximation to [fact] -- a function that returns
correct results for inputs up to [n+1]. Applying [fix] to this
generator returns its _fixed point_, which is a function that
gives the desired behavior for all inputs [n].
(The term "fixed point" is used here in exactly the same sense as
in ordinary mathematics, where a fixed point of a function [f] is
an input [x] such that [f(x) = x]. Here, a fixed point of a
function [F] of type [(Nat->Nat)->(Nat->Nat)] is a function [f] of
type [Nat->Nat] such that [F f] behaves the same as [f].) *)
(** Syntax:
t ::= Terms
| fix t fixed-point operator
| ...
Reduction:
t1 ==> t1'
------------------ (ST_Fix1)
fix t1 ==> fix t1'
-------------------------------------------- (ST_FixAbs)
fix (\xf:T1.t2) ==> [xf:=fix (\xf:T1.t2)] t2
Typing:
Gamma |- t1 : T1->T1
-------------------- (T_Fix)
Gamma |- fix t1 : T1
*)
(** Let's see how [ST_FixAbs] works by reducing [fact 3 = fix F 3],
where
F = (\f. \x. if x=0 then 1 else x * (f (pred x)))
>>
(type annotations are omitted for brevity).
fix F 3
[==>] [ST_FixAbs] + [ST_App1]
(\x. if x=0 then 1 else x * (fix F (pred x))) 3
[==>] [ST_AppAbs]
if 3=0 then 1 else 3 * (fix F (pred 3))
[==>] [ST_If0_Nonzero]
3 * (fix F (pred 3))
[==>] [ST_FixAbs + ST_Mult2]
3 * ((\x. if x=0 then 1 else x * (fix F (pred x))) (pred 3))
[==>] [ST_PredNat + ST_Mult2 + ST_App2]
3 * ((\x. if x=0 then 1 else x * (fix F (pred x))) 2)
[==>] [ST_AppAbs + ST_Mult2]
3 * (if 2=0 then 1 else 2 * (fix F (pred 2)))
[==>] [ST_If0_Nonzero + ST_Mult2]
3 * (2 * (fix F (pred 2)))
[==>] [ST_FixAbs + 2 x ST_Mult2]
3 * (2 * ((\x. if x=0 then 1 else x * (fix F (pred x))) (pred 2)))
[==>] [ST_PredNat + 2 x ST_Mult2 + ST_App2]
3 * (2 * ((\x. if x=0 then 1 else x * (fix F (pred x))) 1))
[==>] [ST_AppAbs + 2 x ST_Mult2]
3 * (2 * (if 1=0 then 1 else 1 * (fix F (pred 1))))
[==>] [ST_If0_Nonzero + 2 x ST_Mult2]
3 * (2 * (1 * (fix F (pred 1))))
[==>] [ST_FixAbs + 3 x ST_Mult2]
3 * (2 * (1 * ((\x. if x=0 then 1 else x * (fix F (pred x))) (pred 1))))
[==>] [ST_PredNat + 3 x ST_Mult2 + ST_App2]
3 * (2 * (1 * ((\x. if x=0 then 1 else x * (fix F (pred x))) 0)))
[==>] [ST_AppAbs + 3 x ST_Mult2]
3 * (2 * (1 * (if 0=0 then 1 else 0 * (fix F (pred 0)))))
[==>] [ST_If0Zero + 3 x ST_Mult2]
3 * (2 * (1 * 1))
[==>] [ST_MultNats + 2 x ST_Mult2]
3 * (2 * 1)
[==>] [ST_MultNats + ST_Mult2]
3 * 2
[==>] [ST_MultNats]
6
*)
(** One important point to note is that, unlike [Fixpoint]
definitions in Coq, there is nothing to prevent functions defined
using [fix] from diverging. *)
(** **** Exercise: 1 star, optional (halve_fix) *)
(** Translate this informal recursive definition into one using [fix]:
halve =
\x:Nat.
if x=0 then 0
else if (pred x)=0 then 0
else 1 + (halve (pred (pred x))))
(* FILL IN HERE *)
[]
*)
(** **** Exercise: 1 star, optional (fact_steps) *)
(** Write down the sequence of steps that the term [fact 1] goes
through to reduce to a normal form (assuming the usual reduction
rules for arithmetic operations).
(* FILL IN HERE *)
[]
*)
(** The ability to form the fixed point of a function of type [T->T]
for any [T] has some surprising consequences. In particular, it
implies that _every_ type is inhabited by some term. To see this,
observe that, for every type [T], we can define the term
fix (\x:T.x)
By [T_Fix] and [T_Abs], this term has type [T]. By [ST_FixAbs]
it reduces to itself, over and over again. Thus it is a
_diverging element_ of [T].
More usefully, here's an example using [fix] to define a
two-argument recursive function:
equal =
fix
(\eq:Nat->Nat->Bool.
\m:Nat. \n:Nat.
if m=0 then iszero n
else if n=0 then false
else eq (pred m) (pred n))
And finally, here is an example where [fix] is used to define a
_pair_ of recursive functions (illustrating the fact that the type
[T1] in the rule [T_Fix] need not be a function type):
evenodd =
fix
(\eo: (Nat->Bool * Nat->Bool).
let e = \n:Nat. if n=0 then true else eo.snd (pred n) in
let o = \n:Nat. if n=0 then false else eo.fst (pred n) in
(e,o))
even = evenodd.fst
odd = evenodd.snd
*)
(* ================================================================= *)
(** ** Records *)
(** As a final example of a basic extension of the STLC, let's look
briefly at how to define _records_ and their types. Intuitively,
records can be obtained from pairs by two straightforward
generalizations: they are n-ary (rather than just binary) and
their fields are accessed by _label_ (rather than position). *)
(** Syntax:
t ::= Terms
| {i1=t1, ..., in=tn} record
| t.i projection
| ...
v ::= Values
| {i1=v1, ..., in=vn} record value
| ...
T ::= Types
| {i1:T1, ..., in:Tn} record type
| ...
The generalization from products should be pretty obvious. But
it's worth noticing the ways in which what we've actually written is
even _more_ informal than the informal syntax we've used in previous
sections and chapters: we've used "[...]" in several places to mean "any
number of these," and we've omitted explicit mention of the usual
side condition that the labels of a record should not contain any
repetitions. *)
(**
Reduction:
ti ==> ti'
------------------------------------ (ST_Rcd)
{i1=v1, ..., im=vm, in=ti, ...}
==> {i1=v1, ..., im=vm, in=ti', ...}
t1 ==> t1'
-------------- (ST_Proj1)
t1.i ==> t1'.i
------------------------- (ST_ProjRcd)
{..., i=vi, ...}.i ==> vi
Again, these rules are a bit informal. For example, the first rule
is intended to be read "if [ti] is the leftmost field that is not a
value and if [ti] steps to [ti'], then the whole record steps..."
In the last rule, the intention is that there should only be one
field called i, and that all the other fields must contain values. *)
(**
The typing rules are also simple:
Gamma |- t1 : T1 ... Gamma |- tn : Tn
-------------------------------------------------- (T_Rcd)
Gamma |- {i1=t1, ..., in=tn} : {i1:T1, ..., in:Tn}
Gamma |- t : {..., i:Ti, ...}
----------------------------- (T_Proj)
Gamma |- t.i : Ti
*)
(** There are several ways to approach formalizing the above
definitions.
- We can directly formalize the syntactic forms and inference
rules, staying as close as possible to the form we've given
them above. This is conceptually straightforward, and it's
probably what we'd want to do if we were building a real
compiler (in particular, it will allow us to print error
messages in the form that programmers will find easy to
understand). But the formal versions of the rules will not be
very pretty or easy to work with, because all the [...]s above
will have to be replaced with explicit quantifications or
comprehensions. For this reason, records are not included in
the extended exercise at the end of this chapter. (It is
still useful to discuss them informally here because they will
help motivate the addition of subtyping to the type system
when we get to the [Sub] chapter.)
- Alternatively, we could look for a smoother way of presenting
records -- for example, a binary presentation with one
constructor for the empty record and another constructor for
adding a single field to an existing record, instead of a
single monolithic constructor that builds a whole record at
once. This is the right way to go if we are primarily
interested in studying the metatheory of the calculi with
records, since it leads to clean and elegant definitions and
proofs. Chapter [Records] shows how this can be done.
- Finally, if we like, we can avoid formalizing records
altogether, by stipulating that record notations are just
informal shorthands for more complex expressions involving
pairs and product types. We sketch this approach in the next
section. *)
(* ----------------------------------------------------------------- *)
(** *** Encoding Records (Optional) *)
(** Let's see how records can be encoded using just pairs and [unit].
First, observe that we can encode arbitrary-size _tuples_ using
nested pairs and the [unit] value. To avoid overloading the pair
notation [(t1,t2)], we'll use curly braces without labels to write
down tuples, so [{}] is the empty tuple, [{5}] is a singleton
tuple, [{5,6}] is a 2-tuple (morally the same as a pair),
[{5,6,7}] is a triple, etc.
{} ----> unit
{t1, t2, ..., tn} ----> (t1, trest)
where {t2, ..., tn} ----> trest
Similarly, we can encode tuple types using nested product types:
{} ----> Unit
{T1, T2, ..., Tn} ----> T1 * TRest
where {T2, ..., Tn} ----> TRest
The operation of projecting a field from a tuple can be encoded
using a sequence of second projections followed by a first projection:
t.0 ----> t.fst
t.(n+1) ----> (t.snd).n
Next, suppose that there is some total ordering on record labels,
so that we can associate each label with a unique natural number.
This number is called the _position_ of the label. For example,
we might assign positions like this:
LABEL POSITION
a 0
b 1
c 2
... ...
bar 1395
... ...
foo 4460
... ...
We use these positions to encode record values as tuples (i.e., as
nested pairs) by sorting the fields according to their positions.
For example:
{a=5, b=6} ----> {5,6}
{a=5, c=7} ----> {5,unit,7}
{c=7, a=5} ----> {5,unit,7}
{c=5, b=3} ----> {unit,3,5}
{f=8,c=5,a=7} ----> {7,unit,5,unit,unit,8}
{f=8,c=5} ----> {unit,unit,5,unit,unit,8}
Note that each field appears in the position associated with its
label, that the size of the tuple is determined by the label with
the highest position, and that we fill in unused positions with
[unit].
We do exactly the same thing with record types:
{a:Nat, b:Nat} ----> {Nat,Nat}
{c:Nat, a:Nat} ----> {Nat,Unit,Nat}
{f:Nat,c:Nat} ----> {Unit,Unit,Nat,Unit,Unit,Nat}
Finally, record projection is encoded as a tuple projection from
the appropriate position:
t.l ----> t.(position of l)
It is not hard to check that all the typing rules for the original
"direct" presentation of records are validated by this
encoding. (The reduction rules are "almost validated" -- not
quite, because the encoding reorders fields.) *)
(** Of course, this encoding will not be very efficient if we
happen to use a record with label [foo]! But things are not
actually as bad as they might seem: for example, if we assume that
our compiler can see the whole program at the same time, we can
_choose_ the numbering of labels so that we assign small positions
to the most frequently used labels. Indeed, there are industrial
compilers that essentially do this! *)
(* ----------------------------------------------------------------- *)
(** *** Variants (Optional) *)
(** Just as products can be generalized to records, sums can be
generalized to n-ary labeled types called _variants_. Instead of
[T1+T2], we can write something like []
where [l1],[l2],... are field labels which are used both to build
instances and as case arm labels.
These n-ary variants give us almost enough mechanism to build
arbitrary inductive data types like lists and trees from
scratch -- the only thing missing is a way to allow _recursion_ in
type definitions. We won't cover this here, but detailed
treatments can be found in many textbooks -- e.g., Types and
Programming Languages [Pierce 2002]. *)
(* ################################################################# *)
(** * Exercise: Formalizing the Extensions *)
(** **** Exercise: 5 stars (STLC_extensions) *)
(** In this exercise, you will formalize some of the extensions
described in this chapter. We've provided the necessary additions
to the syntax of terms and types, and we've included a few
examples that you can test your definitions with to make sure they
are working as expected. You'll fill in the rest of the
definitions and extend all the proofs accordingly.
To get you started, we've provided implementations for:
- numbers
- sums
- lists
- unit
You need to complete the implementations for:
- pairs
- let (which involves binding)
- [fix]
A good strategy is to work on the extensions one at a time, in two
passes, rather than trying to work through the file from start to
finish in a single pass. For each definition or proof, begin by
reading carefully through the parts that are provided for you,
referring to the text in the [Stlc] chapter for high-level
intuitions and the embedded comments for detailed mechanics. *)
Module STLCExtended.
(* ----------------------------------------------------------------- *)
(** *** Syntax *)
Inductive ty : Type :=
| TArrow : ty -> ty -> ty
| TNat : ty
| TUnit : ty
| TProd : ty -> ty -> ty
| TSum : ty -> ty -> ty
| TList : ty -> ty.
Inductive tm : Type :=
(* pure STLC *)
| tvar : id -> tm
| tapp : tm -> tm -> tm
| tabs : id -> ty -> tm -> tm
(* numbers *)
| tnat : nat -> tm
| tsucc : tm -> tm
| tpred : tm -> tm
| tmult : tm -> tm -> tm
| tif0 : tm -> tm -> tm -> tm
(* pairs *)
| tpair : tm -> tm -> tm
| tfst : tm -> tm
| tsnd : tm -> tm
(* units *)
| tunit : tm
(* let *)
| tlet : id -> tm -> tm -> tm
(* i.e., [let x = t1 in t2] *)
(* sums *)
| tinl : ty -> tm -> tm
| tinr : ty -> tm -> tm
| tcase : tm -> id -> tm -> id -> tm -> tm
(* i.e., [case t0 of inl x1 => t1 | inr x2 => t2] *)
(* lists *)
| tnil : ty -> tm
| tcons : tm -> tm -> tm
| tlcase : tm -> tm -> id -> id -> tm -> tm
(* i.e., [lcase t1 of | nil -> t2 | x::y -> t3] *)
(* fix *)
| tfix : tm -> tm.
(** Note that, for brevity, we've omitted booleans and instead
provided a single [if0] form combining a zero test and a
conditional. That is, instead of writing
if x = 0 then ... else ...
we'll write this:
if0 x then ... else ...
*)
(* ----------------------------------------------------------------- *)
(** *** Substitution *)
Fixpoint subst (x:id) (s:tm) (t:tm) : tm :=
match t with
| tvar y =>
if beq_id x y then s else t
| tabs y T t1 =>
tabs y T (if beq_id x y then t1 else (subst x s t1))
| tapp t1 t2 =>
tapp (subst x s t1) (subst x s t2)
| tnat n =>
tnat n
| tsucc t1 =>
tsucc (subst x s t1)
| tpred t1 =>
tpred (subst x s t1)
| tmult t1 t2 =>
tmult (subst x s t1) (subst x s t2)
| tif0 t1 t2 t3 =>
tif0 (subst x s t1) (subst x s t2) (subst x s t3)
(* FILL IN HERE *)
| tunit => tunit
(* FILL IN HERE *)
| tinl T t1 =>
tinl T (subst x s t1)
| tinr T t1 =>
tinr T (subst x s t1)
| tcase t0 y1 t1 y2 t2 =>
tcase (subst x s t0)
y1 (if beq_id x y1 then t1 else (subst x s t1))
y2 (if beq_id x y2 then t2 else (subst x s t2))
| tnil T =>
tnil T
| tcons t1 t2 =>
tcons (subst x s t1) (subst x s t2)
| tlcase t1 t2 y1 y2 t3 =>
tlcase (subst x s t1) (subst x s t2) y1 y2
(if beq_id x y1 then
t3
else if beq_id x y2 then t3
else (subst x s t3))
(* FILL IN HERE *)
| _ => t (* ... and delete this line *)
end.
Notation "'[' x ':=' s ']' t" := (subst x s t) (at level 20).
(* ----------------------------------------------------------------- *)
(** *** Reduction *)
(** Next we define the values of our language. *)
Inductive value : tm -> Prop :=
| v_abs : forall x T11 t12,
value (tabs x T11 t12)
(* Numbers are values: *)
| v_nat : forall n1,
value (tnat n1)
(* A pair is a value if both components are: *)
| v_pair : forall v1 v2,
value v1 ->
value v2 ->
value (tpair v1 v2)
(* A unit is always a value *)
| v_unit : value tunit
(* A tagged value is a value: *)
| v_inl : forall v T,
value v ->
value (tinl T v)
| v_inr : forall v T,
value v ->
value (tinr T v)
(* A list is a value iff its head and tail are values: *)
| v_lnil : forall T, value (tnil T)
| v_lcons : forall v1 vl,
value v1 ->
value vl ->
value (tcons v1 vl)
.
Hint Constructors value.
Reserved Notation "t1 '==>' t2" (at level 40).
Inductive step : tm -> tm -> Prop :=
| ST_AppAbs : forall x T11 t12 v2,
value v2 ->
(tapp (tabs x T11 t12) v2) ==> [x:=v2]t12
| ST_App1 : forall t1 t1' t2,
t1 ==> t1' ->
(tapp t1 t2) ==> (tapp t1' t2)
| ST_App2 : forall v1 t2 t2',
value v1 ->
t2 ==> t2' ->
(tapp v1 t2) ==> (tapp v1 t2')
(* nats *)
| ST_Succ1 : forall t1 t1',
t1 ==> t1' ->
(tsucc t1) ==> (tsucc t1')
| ST_SuccNat : forall n1,
(tsucc (tnat n1)) ==> (tnat (S n1))
| ST_Pred : forall t1 t1',
t1 ==> t1' ->
(tpred t1) ==> (tpred t1')
| ST_PredNat : forall n1,
(tpred (tnat n1)) ==> (tnat (pred n1))
| ST_Mult1 : forall t1 t1' t2,
t1 ==> t1' ->
(tmult t1 t2) ==> (tmult t1' t2)
| ST_Mult2 : forall v1 t2 t2',
value v1 ->
t2 ==> t2' ->
(tmult v1 t2) ==> (tmult v1 t2')
| ST_MultNats : forall n1 n2,
(tmult (tnat n1) (tnat n2)) ==> (tnat (mult n1 n2))
| ST_If01 : forall t1 t1' t2 t3,
t1 ==> t1' ->
(tif0 t1 t2 t3) ==> (tif0 t1' t2 t3)
| ST_If0Zero : forall t2 t3,
(tif0 (tnat 0) t2 t3) ==> t2
| ST_If0Nonzero : forall n t2 t3,
(tif0 (tnat (S n)) t2 t3) ==> t3
(* pairs *)
(* FILL IN HERE *)
(* let *)
(* FILL IN HERE *)
(* sums *)
| ST_Inl : forall t1 t1' T,
t1 ==> t1' ->
(tinl T t1) ==> (tinl T t1')
| ST_Inr : forall t1 t1' T,
t1 ==> t1' ->
(tinr T t1) ==> (tinr T t1')
| ST_Case : forall t0 t0' x1 t1 x2 t2,
t0 ==> t0' ->
(tcase t0 x1 t1 x2 t2) ==> (tcase t0' x1 t1 x2 t2)
| ST_CaseInl : forall v0 x1 t1 x2 t2 T,
value v0 ->
(tcase (tinl T v0) x1 t1 x2 t2) ==> [x1:=v0]t1
| ST_CaseInr : forall v0 x1 t1 x2 t2 T,
value v0 ->
(tcase (tinr T v0) x1 t1 x2 t2) ==> [x2:=v0]t2
(* lists *)
| ST_Cons1 : forall t1 t1' t2,
t1 ==> t1' ->
(tcons t1 t2) ==> (tcons t1' t2)
| ST_Cons2 : forall v1 t2 t2',
value v1 ->
t2 ==> t2' ->
(tcons v1 t2) ==> (tcons v1 t2')
| ST_Lcase1 : forall t1 t1' t2 x1 x2 t3,
t1 ==> t1' ->
(tlcase t1 t2 x1 x2 t3) ==> (tlcase t1' t2 x1 x2 t3)
| ST_LcaseNil : forall T t2 x1 x2 t3,
(tlcase (tnil T) t2 x1 x2 t3) ==> t2
| ST_LcaseCons : forall v1 vl t2 x1 x2 t3,
value v1 ->
value vl ->
(tlcase (tcons v1 vl) t2 x1 x2 t3) ==> (subst x2 vl (subst x1 v1 t3))
(* fix *)
(* FILL IN HERE *)
where "t1 '==>' t2" := (step t1 t2).
Notation multistep := (multi step).
Notation "t1 '==>*' t2" := (multistep t1 t2) (at level 40).
Hint Constructors step.
(* ----------------------------------------------------------------- *)
(** *** Typing *)
Definition context := partial_map ty.
(** Next we define the typing rules. These are nearly direct
transcriptions of the inference rules shown above. *)
Reserved Notation "Gamma '|-' t '\in' T" (at level 40).
Inductive has_type : context -> tm -> ty -> Prop :=
(* Typing rules for proper terms *)
| T_Var : forall Gamma x T,
Gamma x = Some T ->
Gamma |- (tvar x) \in T
| T_Abs : forall Gamma x T11 T12 t12,
(update Gamma x T11) |- t12 \in T12 ->
Gamma |- (tabs x T11 t12) \in (TArrow T11 T12)
| T_App : forall T1 T2 Gamma t1 t2,
Gamma |- t1 \in (TArrow T1 T2) ->
Gamma |- t2 \in T1 ->
Gamma |- (tapp t1 t2) \in T2
(* nats *)
| T_Nat : forall Gamma n1,
Gamma |- (tnat n1) \in TNat
| T_Succ : forall Gamma t1,
Gamma |- t1 \in TNat ->
Gamma |- (tsucc t1) \in TNat
| T_Pred : forall Gamma t1,
Gamma |- t1 \in TNat ->
Gamma |- (tpred t1) \in TNat
| T_Mult : forall Gamma t1 t2,
Gamma |- t1 \in TNat ->
Gamma |- t2 \in TNat ->
Gamma |- (tmult t1 t2) \in TNat
| T_If0 : forall Gamma t1 t2 t3 T1,
Gamma |- t1 \in TNat ->
Gamma |- t2 \in T1 ->
Gamma |- t3 \in T1 ->
Gamma |- (tif0 t1 t2 t3) \in T1
(* pairs *)
(* FILL IN HERE *)
(* unit *)
| T_Unit : forall Gamma,
Gamma |- tunit \in TUnit
(* let *)
(* FILL IN HERE *)
(* sums *)
| T_Inl : forall Gamma t1 T1 T2,
Gamma |- t1 \in T1 ->
Gamma |- (tinl T2 t1) \in (TSum T1 T2)
| T_Inr : forall Gamma t2 T1 T2,
Gamma |- t2 \in T2 ->
Gamma |- (tinr T1 t2) \in (TSum T1 T2)
| T_Case : forall Gamma t0 x1 T1 t1 x2 T2 t2 T,
Gamma |- t0 \in (TSum T1 T2) ->
(update Gamma x1 T1) |- t1 \in T ->
(update Gamma x2 T2) |- t2 \in T ->
Gamma |- (tcase t0 x1 t1 x2 t2) \in T
(* lists *)
| T_Nil : forall Gamma T,
Gamma |- (tnil T) \in (TList T)
| T_Cons : forall Gamma t1 t2 T1,
Gamma |- t1 \in T1 ->
Gamma |- t2 \in (TList T1) ->
Gamma |- (tcons t1 t2) \in (TList T1)
| T_Lcase : forall Gamma t1 T1 t2 x1 x2 t3 T2,
Gamma |- t1 \in (TList T1) ->
Gamma |- t2 \in T2 ->
(update (update Gamma x2 (TList T1)) x1 T1) |- t3 \in T2 ->
Gamma |- (tlcase t1 t2 x1 x2 t3) \in T2
(* fix *)
(* FILL IN HERE *)
where "Gamma '|-' t '\in' T" := (has_type Gamma t T).
Hint Constructors has_type.
(* ================================================================= *)
(** ** Examples *)
(** This section presents formalized versions of the examples from
above (plus several more). The ones at the beginning focus on
specific features; you can use these to make sure your definition
of a given feature is reasonable before moving on to extending the
proofs later in the file with the cases relating to this feature.
The later examples require all the features together, so you'll
need to come back to these when you've got all the definitions
filled in. *)
Module Examples.
(* ----------------------------------------------------------------- *)
(** *** Preliminaries *)
(** First, let's define a few variable names: *)
Notation x := (Id "x").
Notation y := (Id "y").
Notation a := (Id "a").
Notation f := (Id "f").
Notation g := (Id "g").
Notation l := (Id "l").
Notation k := (Id "k").
Notation i1 := (Id "i1").
Notation i2 := (Id "i2").
Notation processSum := (Id "processSum").
Notation n := (Id "n").
Notation eq := (Id "eq").
Notation m := (Id "m").
Notation evenodd := (Id "evenodd").
Notation even := (Id "even").
Notation odd := (Id "odd").
Notation eo := (Id "eo").
(** Next, a bit of Coq hackery to automate searching for typing
derivations. You don't need to understand this bit in detail --
just have a look over it so that you'll know what to look for if
you ever find yourself needing to make custom extensions to
[auto].
The following [Hint] declarations say that, whenever [auto]
arrives at a goal of the form [(Gamma |- (tapp e1 e1) \in T)], it
should consider [eapply T_App], leaving an existential variable
for the middle type T1, and similar for [lcase]. That variable
will then be filled in during the search for type derivations for
[e1] and [e2]. We also include a hint to "try harder" when
solving equality goals; this is useful to automate uses of
[T_Var] (which includes an equality as a precondition). *)
Hint Extern 2 (has_type _ (tapp _ _) _) =>
eapply T_App; auto.
Hint Extern 2 (has_type _ (tlcase _ _ _ _ _) _) =>
eapply T_Lcase; auto.
Hint Extern 2 (_ = _) => compute; reflexivity.
(* ----------------------------------------------------------------- *)
(** *** Numbers *)
Module Numtest.
(* if0 (pred (succ (pred (2 * 0))) then 5 else 6 *)
Definition test :=
tif0
(tpred
(tsucc
(tpred
(tmult
(tnat 2)
(tnat 0)))))
(tnat 5)
(tnat 6).
(** Remove the comment braces once you've implemented enough of the
definitions that you think this should work. *)
(*
Example typechecks :
empty |- test \in TNat.
Proof.
unfold test.
(* This typing derivation is quite deep, so we need
to increase the max search depth of [auto] from the
default 5 to 10. *)
auto 10.
Qed.
Example numtest_reduces :
test ==>* tnat 5.
Proof.
unfold test. normalize.
Qed.
*)
End Numtest.
(* ----------------------------------------------------------------- *)
(** *** Products *)
Module Prodtest.
(* ((5,6),7).fst.snd *)
Definition test :=
tsnd
(tfst
(tpair
(tpair
(tnat 5)
(tnat 6))
(tnat 7))).
(*
Example typechecks :
empty |- test \in TNat.
Proof. unfold test. eauto 15. Qed.
Example reduces :
test ==>* tnat 6.
Proof. unfold test. normalize. Qed.
*)
End Prodtest.
(* ----------------------------------------------------------------- *)
(** *** [let] *)
Module LetTest.
(* let x = pred 6 in succ x *)
Definition test :=
tlet
x
(tpred (tnat 6))
(tsucc (tvar x)).
(*
Example typechecks :
empty |- test \in TNat.
Proof. unfold test. eauto 15. Qed.
Example reduces :
test ==>* tnat 6.
Proof. unfold test. normalize. Qed.
*)
End LetTest.
(* ----------------------------------------------------------------- *)
(** *** Sums *)
Module Sumtest1.
(* case (inl Nat 5) of
inl x => x
| inr y => y *)
Definition test :=
tcase (tinl TNat (tnat 5))
x (tvar x)
y (tvar y).
(*
Example typechecks :
empty |- test \in TNat.
Proof. unfold test. eauto 15. Qed.
Example reduces :
test ==>* (tnat 5).
Proof. unfold test. normalize. Qed.
*)
End Sumtest1.
Module Sumtest2.
(* let processSum =
\x:Nat+Nat.
case x of
inl n => n
inr n => if0 n then 1 else 0 in
(processSum (inl Nat 5), processSum (inr Nat 5)) *)
Definition test :=
tlet
processSum
(tabs x (TSum TNat TNat)
(tcase (tvar x)
n (tvar n)
n (tif0 (tvar n) (tnat 1) (tnat 0))))
(tpair
(tapp (tvar processSum) (tinl TNat (tnat 5)))
(tapp (tvar processSum) (tinr TNat (tnat 5)))).
(*
Example typechecks :
empty |- test \in (TProd TNat TNat).
Proof. unfold test. eauto 15. Qed.
Example reduces :
test ==>* (tpair (tnat 5) (tnat 0)).
Proof. unfold test. normalize. Qed.
*)
End Sumtest2.
(* ----------------------------------------------------------------- *)
(** *** Lists *)
Module ListTest.
(* let l = cons 5 (cons 6 (nil Nat)) in
lcase l of
nil => 0
| x::y => x*x *)
Definition test :=
tlet l
(tcons (tnat 5) (tcons (tnat 6) (tnil TNat)))
(tlcase (tvar l)
(tnat 0)
x y (tmult (tvar x) (tvar x))).
(*
Example typechecks :
empty |- test \in TNat.
Proof. unfold test. eauto 20. Qed.
Example reduces :
test ==>* (tnat 25).
Proof. unfold test. normalize. Qed.
*)
End ListTest.
(* ----------------------------------------------------------------- *)
(** *** [fix] *)
Module FixTest1.
(* fact := fix
(\f:nat->nat.
\a:nat.
if a=0 then 1 else a * (f (pred a))) *)
Definition fact :=
tfix
(tabs f (TArrow TNat TNat)
(tabs a TNat
(tif0
(tvar a)
(tnat 1)
(tmult
(tvar a)
(tapp (tvar f) (tpred (tvar a))))))).
(** (Warning: you may be able to typecheck [fact] but still have some
rules wrong!) *)
(*
Example fact_typechecks :
empty |- fact \in (TArrow TNat TNat).
Proof. unfold fact. auto 10.
Qed.
*)
(*
Example fact_example:
(tapp fact (tnat 4)) ==>* (tnat 24).
Proof. unfold fact. normalize. Qed.
*)
End FixTest1.
Module FixTest2.
(* map :=
\g:nat->nat.
fix
(\f:[nat]->[nat].
\l:[nat].
case l of
| [] -> []
| x::l -> (g x)::(f l)) *)
Definition map :=
tabs g (TArrow TNat TNat)
(tfix
(tabs f (TArrow (TList TNat) (TList TNat))
(tabs l (TList TNat)
(tlcase (tvar l)
(tnil TNat)
a l (tcons (tapp (tvar g) (tvar a))
(tapp (tvar f) (tvar l))))))).
(*
(* Make sure you've uncommented the last [Hint Extern] above... *)
Example map_typechecks :
empty |- map \in
(TArrow (TArrow TNat TNat)
(TArrow (TList TNat)
(TList TNat))).
Proof. unfold map. auto 10. Qed.
Example map_example :
tapp (tapp map (tabs a TNat (tsucc (tvar a))))
(tcons (tnat 1) (tcons (tnat 2) (tnil TNat)))
==>* (tcons (tnat 2) (tcons (tnat 3) (tnil TNat))).
Proof. unfold map. normalize. Qed.
*)
End FixTest2.
Module FixTest3.
(* equal =
fix
(\eq:Nat->Nat->Bool.
\m:Nat. \n:Nat.
if0 m then (if0 n then 1 else 0)
else if0 n then 0
else eq (pred m) (pred n)) *)
Definition equal :=
tfix
(tabs eq (TArrow TNat (TArrow TNat TNat))
(tabs m TNat
(tabs n TNat
(tif0 (tvar m)
(tif0 (tvar n) (tnat 1) (tnat 0))
(tif0 (tvar n)
(tnat 0)
(tapp (tapp (tvar eq)
(tpred (tvar m)))
(tpred (tvar n)))))))).
(*
Example equal_typechecks :
empty |- equal \in (TArrow TNat (TArrow TNat TNat)).
Proof. unfold equal. auto 10.
Qed.
*)
(*
Example equal_example1:
(tapp (tapp equal (tnat 4)) (tnat 4)) ==>* (tnat 1).
Proof. unfold equal. normalize. Qed.
*)
(*
Example equal_example2:
(tapp (tapp equal (tnat 4)) (tnat 5)) ==>* (tnat 0).
Proof. unfold equal. normalize. Qed.
*)
End FixTest3.
Module FixTest4.
(* let evenodd =
fix
(\eo: (Nat->Nat * Nat->Nat).
let e = \n:Nat. if0 n then 1 else eo.snd (pred n) in
let o = \n:Nat. if0 n then 0 else eo.fst (pred n) in
(e,o)) in
let even = evenodd.fst in
let odd = evenodd.snd in
(even 3, even 4)
*)
Definition eotest :=
tlet evenodd
(tfix
(tabs eo (TProd (TArrow TNat TNat) (TArrow TNat TNat))
(tpair
(tabs n TNat
(tif0 (tvar n)
(tnat 1)
(tapp (tsnd (tvar eo)) (tpred (tvar n)))))
(tabs n TNat
(tif0 (tvar n)
(tnat 0)
(tapp (tfst (tvar eo)) (tpred (tvar n))))))))
(tlet even (tfst (tvar evenodd))
(tlet odd (tsnd (tvar evenodd))
(tpair
(tapp (tvar even) (tnat 3))
(tapp (tvar even) (tnat 4))))).
(*
Example eotest_typechecks :
empty |- eotest \in (TProd TNat TNat).
Proof. unfold eotest. eauto 30.
Qed.
*)
(*
Example eotest_example1:
eotest ==>* (tpair (tnat 0) (tnat 1)).
Proof. unfold eotest. normalize. Qed.
*)
End FixTest4.
End Examples.
(* ================================================================= *)
(** ** Properties of Typing *)
(** The proofs of progress and preservation for this enriched system
are essentially the same (though of course longer) as for the pure
STLC. *)
(* ----------------------------------------------------------------- *)
(** *** Progress *)
Theorem progress : forall t T,
empty |- t \in T ->
value t \/ exists t', t ==> t'.
Proof with eauto.
(* Theorem: Suppose empty |- t : T. Then either
1. t is a value, or
2. t ==> t' for some t'.
Proof: By induction on the given typing derivation. *)
intros t T Ht.
remember empty as Gamma.
generalize dependent HeqGamma.
induction Ht; intros HeqGamma; subst.
- (* T_Var *)
(* The final rule in the given typing derivation cannot be
[T_Var], since it can never be the case that
[empty |- x : T] (since the context is empty). *)
inversion H.
- (* T_Abs *)
(* If the [T_Abs] rule was the last used, then
[t = tabs x T11 t12], which is a value. *)
left...
- (* T_App *)
(* If the last rule applied was T_App, then [t = t1 t2],
and we know from the form of the rule that
[empty |- t1 : T1 -> T2]
[empty |- t2 : T1]
By the induction hypothesis, each of t1 and t2 either is
a value or can take a step. *)
right.
destruct IHHt1; subst...
+ (* t1 is a value *)
destruct IHHt2; subst...
* (* t2 is a value *)
(* If both [t1] and [t2] are values, then we know that
[t1 = tabs x T11 t12], since abstractions are the
only values that can have an arrow type. But
[(tabs x T11 t12) t2 ==> [x:=t2]t12] by [ST_AppAbs]. *)
inversion H; subst; try solve_by_invert.
exists (subst x t2 t12)...
* (* t2 steps *)
(* If [t1] is a value and [t2 ==> t2'],
then [t1 t2 ==> t1 t2'] by [ST_App2]. *)
inversion H0 as [t2' Hstp]. exists (tapp t1 t2')...
+ (* t1 steps *)
(* Finally, If [t1 ==> t1'], then [t1 t2 ==> t1' t2]
by [ST_App1]. *)
inversion H as [t1' Hstp]. exists (tapp t1' t2)...
- (* T_Nat *)
left...
- (* T_Succ *)
right.
destruct IHHt...
+ (* t1 is a value *)
inversion H; subst; try solve_by_invert.
exists (tnat (S n1))...
+ (* t1 steps *)
inversion H as [t1' Hstp].
exists (tsucc t1')...
- (* T_Pred *)
right.
destruct IHHt...
+ (* t1 is a value *)
inversion H; subst; try solve_by_invert.
exists (tnat (pred n1))...
+ (* t1 steps *)
inversion H as [t1' Hstp].
exists (tpred t1')...
- (* T_Mult *)
right.
destruct IHHt1...
+ (* t1 is a value *)
destruct IHHt2...
* (* t2 is a value *)
inversion H; subst; try solve_by_invert.
inversion H0; subst; try solve_by_invert.
exists (tnat (mult n1 n0))...
* (* t2 steps *)
inversion H0 as [t2' Hstp].
exists (tmult t1 t2')...
+ (* t1 steps *)
inversion H as [t1' Hstp].
exists (tmult t1' t2)...
- (* T_If0 *)
right.
destruct IHHt1...
+ (* t1 is a value *)
inversion H; subst; try solve_by_invert.
destruct n1 as [|n1'].
* (* n1=0 *)
exists t2...
* (* n1<>0 *)
exists t3...
+ (* t1 steps *)
inversion H as [t1' H0].
exists (tif0 t1' t2 t3)...
(* FILL IN HERE *)
- (* T_Unit *)
left...
(* let *)
(* FILL IN HERE *)
- (* T_Inl *)
destruct IHHt...
+ (* t1 steps *)
right. inversion H as [t1' Hstp]...
(* exists (tinl _ t1')... *)
- (* T_Inr *)
destruct IHHt...
+ (* t1 steps *)
right. inversion H as [t1' Hstp]...
(* exists (tinr _ t1')... *)
- (* T_Case *)
right.
destruct IHHt1...
+ (* t0 is a value *)
inversion H; subst; try solve_by_invert.
* (* t0 is inl *)
exists ([x1:=v]t1)...
* (* t0 is inr *)
exists ([x2:=v]t2)...
+ (* t0 steps *)
inversion H as [t0' Hstp].
exists (tcase t0' x1 t1 x2 t2)...
- (* T_Nil *)
left...
- (* T_Cons *)
destruct IHHt1...
+ (* head is a value *)
destruct IHHt2...
* (* tail steps *)
right. inversion H0 as [t2' Hstp].
exists (tcons t1 t2')...
+ (* head steps *)
right. inversion H as [t1' Hstp].
exists (tcons t1' t2)...
- (* T_Lcase *)
right.
destruct IHHt1...
+ (* t1 is a value *)
inversion H; subst; try solve_by_invert.
* (* t1=tnil *)
exists t2...
* (* t1=tcons v1 vl *)
exists ([x2:=vl]([x1:=v1]t3))...
+ (* t1 steps *)
inversion H as [t1' Hstp].
exists (tlcase t1' t2 x1 x2 t3)...
(* fix *)
(* FILL IN HERE *)
Qed.
(* ----------------------------------------------------------------- *)
(** *** Context Invariance *)
Inductive appears_free_in : id -> tm -> Prop :=
| afi_var : forall x,
appears_free_in x (tvar x)
| afi_app1 : forall x t1 t2,
appears_free_in x t1 -> appears_free_in x (tapp t1 t2)
| afi_app2 : forall x t1 t2,
appears_free_in x t2 -> appears_free_in x (tapp t1 t2)
| afi_abs : forall x y T11 t12,
y <> x ->
appears_free_in x t12 ->
appears_free_in x (tabs y T11 t12)
(* nats *)
| afi_succ : forall x t,
appears_free_in x t ->
appears_free_in x (tsucc t)
| afi_pred : forall x t,
appears_free_in x t ->
appears_free_in x (tpred t)
| afi_mult1 : forall x t1 t2,
appears_free_in x t1 ->
appears_free_in x (tmult t1 t2)
| afi_mult2 : forall x t1 t2,
appears_free_in x t2 ->
appears_free_in x (tmult t1 t2)
| afi_if01 : forall x t1 t2 t3,
appears_free_in x t1 ->
appears_free_in x (tif0 t1 t2 t3)
| afi_if02 : forall x t1 t2 t3,
appears_free_in x t2 ->
appears_free_in x (tif0 t1 t2 t3)
| afi_if03 : forall x t1 t2 t3,
appears_free_in x t3 ->
appears_free_in x (tif0 t1 t2 t3)
(* pairs *)
(* FILL IN HERE *)
(* let *)
(* FILL IN HERE *)
(* sums *)
| afi_inl : forall x t T,
appears_free_in x t ->
appears_free_in x (tinl T t)
| afi_inr : forall x t T,
appears_free_in x t ->
appears_free_in x (tinr T t)
| afi_case0 : forall x t0 x1 t1 x2 t2,
appears_free_in x t0 ->
appears_free_in x (tcase t0 x1 t1 x2 t2)
| afi_case1 : forall x t0 x1 t1 x2 t2,
x1 <> x ->
appears_free_in x t1 ->
appears_free_in x (tcase t0 x1 t1 x2 t2)
| afi_case2 : forall x t0 x1 t1 x2 t2,
x2 <> x ->
appears_free_in x t2 ->
appears_free_in x (tcase t0 x1 t1 x2 t2)
(* lists *)
| afi_cons1 : forall x t1 t2,
appears_free_in x t1 ->
appears_free_in x (tcons t1 t2)
| afi_cons2 : forall x t1 t2,
appears_free_in x t2 ->
appears_free_in x (tcons t1 t2)
| afi_lcase1 : forall x t1 t2 y1 y2 t3,
appears_free_in x t1 ->
appears_free_in x (tlcase t1 t2 y1 y2 t3)
| afi_lcase2 : forall x t1 t2 y1 y2 t3,
appears_free_in x t2 ->
appears_free_in x (tlcase t1 t2 y1 y2 t3)
| afi_lcase3 : forall x t1 t2 y1 y2 t3,
y1 <> x ->
y2 <> x ->
appears_free_in x t3 ->
appears_free_in x (tlcase t1 t2 y1 y2 t3)
(* fix *)
(* FILL IN HERE *)
.
Hint Constructors appears_free_in.
Lemma context_invariance : forall Gamma Gamma' t S,
Gamma |- t \in S ->
(forall x, appears_free_in x t -> Gamma x = Gamma' x) ->
Gamma' |- t \in S.
Proof with eauto.
intros. generalize dependent Gamma'.
induction H;
intros Gamma' Heqv...
- (* T_Var *)
apply T_Var... rewrite <- Heqv...
- (* T_Abs *)
apply T_Abs... apply IHhas_type. intros y Hafi.
unfold update, t_update.
destruct (beq_idP x y)...
- (* T_Mult *)
apply T_Mult...
- (* T_If0 *)
apply T_If0...
(* pair *)
(* FILL IN HERE *)
(* let *)
(* FILL IN HERE *)
- (* T_Case *)
eapply T_Case...
+ apply IHhas_type2. intros y Hafi.
unfold update, t_update.
destruct (beq_idP x1 y)...
+ apply IHhas_type3. intros y Hafi.
unfold update, t_update.
destruct (beq_idP x2 y)...
- (* T_Cons *)
apply T_Cons...
- (* T_Lcase *)
eapply T_Lcase... apply IHhas_type3. intros y Hafi.
unfold update, t_update.
destruct (beq_idP x1 y)...
destruct (beq_idP x2 y)...
Qed.
Lemma free_in_context : forall x t T Gamma,
appears_free_in x t ->
Gamma |- t \in T ->
exists T', Gamma x = Some T'.
Proof with eauto.
intros x t T Gamma Hafi Htyp.
induction Htyp; inversion Hafi; subst...
- (* T_Abs *)
destruct IHHtyp as [T' Hctx]... exists T'.
unfold update, t_update in Hctx.
rewrite false_beq_id in Hctx...
(* let *)
(* FILL IN HERE *)
(* T_Case *)
- (* left *)
destruct IHHtyp2 as [T' Hctx]... exists T'.
unfold update, t_update in Hctx.
rewrite false_beq_id in Hctx...
- (* right *)
destruct IHHtyp3 as [T' Hctx]... exists T'.
unfold update, t_update in Hctx.
rewrite false_beq_id in Hctx...
- (* T_Lcase *)
clear Htyp1 IHHtyp1 Htyp2 IHHtyp2.
destruct IHHtyp3 as [T' Hctx]... exists T'.
unfold update, t_update in Hctx.
rewrite false_beq_id in Hctx...
rewrite false_beq_id in Hctx...
Qed.
(* ----------------------------------------------------------------- *)
(** *** Substitution *)
Lemma substitution_preserves_typing : forall Gamma x U v t S,
(update Gamma x U) |- t \in S ->
empty |- v \in U ->
Gamma |- ([x:=v]t) \in S.
Proof with eauto.
(* Theorem: If Gamma,x:U |- t : S and empty |- v : U, then
Gamma |- [x:=v]t : S. *)
intros Gamma x U v t S Htypt Htypv.
generalize dependent Gamma. generalize dependent S.
(* Proof: By induction on the term t. Most cases follow
directly from the IH, with the exception of tvar
and tabs. These aren't automatic because we must
reason about how the variables interact. *)
induction t;
intros S Gamma Htypt; simpl; inversion Htypt; subst...
- (* tvar *)
simpl. rename i into y.
(* If t = y, we know that
[empty |- v : U] and
[Gamma,x:U |- y : S]
and, by inversion, [update Gamma x U y = Some S]. We want to
show that [Gamma |- [x:=v]y : S].
There are two cases to consider: either [x=y] or [x<>y]. *)
unfold update, t_update in H1.
destruct (beq_idP x y).
+ (* x=y *)
(* If [x = y], then we know that [U = S], and that
[[x:=v]y = v]. So what we really must show is
that if [empty |- v : U] then [Gamma |- v : U].
We have already proven a more general version
of this theorem, called context invariance. *)
subst.
inversion H1; subst. clear H1.
eapply context_invariance...
intros x Hcontra.
destruct (free_in_context _ _ S empty Hcontra)
as [T' HT']...
inversion HT'.
+ (* x<>y *)
(* If [x <> y], then [Gamma y = Some S] and the substitution has no
effect. We can show that [Gamma |- y : S] by [T_Var]. *)
apply T_Var...
- (* tabs *)
rename i into y. rename t into T11.
(* If [t = tabs y T11 t0], then we know that
[Gamma,x:U |- tabs y T11 t0 : T11->T12]
[Gamma,x:U,y:T11 |- t0 : T12]
[empty |- v : U]
As our IH, we know that forall S Gamma,
[Gamma,x:U |- t0 : S -> Gamma |- [x:=v]t0 : S].
We can calculate that
[x:=v]t = tabs y T11 (if beq_id x y then t0 else [x:=v]t0)
And we must show that [Gamma |- [x:=v]t : T11->T12]. We know
we will do so using [T_Abs], so it remains to be shown that:
[Gamma,y:T11 |- if beq_id x y then t0 else [x:=v]t0 : T12]
We consider two cases: [x = y] and [x <> y].
*)
apply T_Abs...
destruct (beq_idP x y) as [Hxy|Hxy].
+ (* x=y *)
(* If [x = y], then the substitution has no effect. Context
invariance shows that [Gamma,y:U,y:T11] and [Gamma,y:T11] are
equivalent. Since the former context shows that
[t0 : T12], so does the latter. *)
eapply context_invariance...
subst.
intros x Hafi. unfold update, t_update.
destruct (beq_id y x)...
+ (* x<>y *)
(* If [x <> y], then the IH and context invariance allow
us to show that
[Gamma,x:U,y:T11 |- t0 : T12] =>
[Gamma,y:T11,x:U |- t0 : T12] =>
[Gamma,y:T11 |- [x:=v]t0 : T12] *)
apply IHt. eapply context_invariance...
intros z Hafi. unfold update, t_update.
destruct (beq_idP y z) as [Hyz|Hyz]...
subst.
rewrite false_beq_id...
(* let *)
(* FILL IN HERE *)
- (* tcase *)
rename i into x1. rename i0 into x2.
eapply T_Case...
+ (* left arm *)
destruct (beq_idP x x1) as [Hxx1|Hxx1].
* (* x = x1 *)
eapply context_invariance...
subst.
intros z Hafi. unfold update, t_update.
destruct (beq_id x1 z)...
* (* x <> x1 *)
apply IHt2. eapply context_invariance...
intros z Hafi. unfold update, t_update.
destruct (beq_idP x1 z) as [Hx1z|Hx1z]...
subst. rewrite false_beq_id...
+ (* right arm *)
destruct (beq_idP x x2) as [Hxx2|Hxx2].
* (* x = x2 *)
eapply context_invariance...
subst.
intros z Hafi. unfold update, t_update.
destruct (beq_id x2 z)...
* (* x <> x2 *)
apply IHt3. eapply context_invariance...
intros z Hafi. unfold update, t_update.
destruct (beq_idP x2 z)...
subst. rewrite false_beq_id...
- (* tlcase *)
rename i into y1. rename i0 into y2.
eapply T_Lcase...
destruct (beq_idP x y1).
+ (* x=y1 *)
simpl.
eapply context_invariance...
subst.
intros z Hafi. unfold update, t_update.
destruct (beq_idP y1 z)...
+ (* x<>y1 *)
destruct (beq_idP x y2).
* (* x=y2 *)
eapply context_invariance...
subst.
intros z Hafi. unfold update, t_update.
destruct (beq_idP y2 z)...
* (* x<>y2 *)
apply IHt3. eapply context_invariance...
intros z Hafi. unfold update, t_update.
destruct (beq_idP y1 z)...
subst. rewrite false_beq_id...
destruct (beq_idP y2 z)...
subst. rewrite false_beq_id...
Qed.
(* ----------------------------------------------------------------- *)
(** *** Preservation *)
Theorem preservation : forall t t' T,
empty |- t \in T ->
t ==> t' ->
empty |- t' \in T.
Proof with eauto.
intros t t' T HT.
(* Theorem: If [empty |- t : T] and [t ==> t'], then
[empty |- t' : T]. *)
remember empty as Gamma. generalize dependent HeqGamma.
generalize dependent t'.
(* Proof: By induction on the given typing derivation. Many
cases are contradictory ([T_Var], [T_Abs]). We show just
the interesting ones. *)
induction HT;
intros t' HeqGamma HE; subst; inversion HE; subst...
- (* T_App *)
(* If the last rule used was [T_App], then [t = t1 t2], and
three rules could have been used to show [t ==> t']:
[ST_App1], [ST_App2], and [ST_AppAbs]. In the first two
cases, the result follows directly from the IH. *)
inversion HE; subst...
+ (* ST_AppAbs *)
(* For the third case, suppose
[t1 = tabs x T11 t12]
and
[t2 = v2].
We must show that [empty |- [x:=v2]t12 : T2].
We know by assumption that
[empty |- tabs x T11 t12 : T1->T2]
and by inversion
[x:T1 |- t12 : T2]
We have already proven that substitution preserves
typing, and
[empty |- v2 : T1]
by assumption, so we are done. *)
apply substitution_preserves_typing with T1...
inversion HT1...
(* fst and snd *)
(* FILL IN HERE *)
(* let *)
(* FILL IN HERE *)
(* T_Case *)
- (* ST_CaseInl *)
inversion HT1; subst.
eapply substitution_preserves_typing...
- (* ST_CaseInr *)
inversion HT1; subst.
eapply substitution_preserves_typing...
- (* T_Lcase *)
+ (* ST_LcaseCons *)
inversion HT1; subst.
apply substitution_preserves_typing with (TList T1)...
apply substitution_preserves_typing with T1...
(* fix *)
(* FILL IN HERE *)
Qed.
(** [] *)
End STLCExtended.
(* $Date: 2016-12-02 15:58:09 -0500 (Fri, 02 Dec 2016) $ *)