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Undo some overzealous vspace tweaks
author Adam Chlipala <adam@chlipala.net>
date Wed, 01 Aug 2012 16:09:37 -0400
parents f923024bd284
children f28bdd8414e0
rev   line source
adam@398 1 (* Copyright (c) 2008-2012, Adam Chlipala
adamc@62 2 *
adamc@62 3 * This work is licensed under a
adamc@62 4 * Creative Commons Attribution-Noncommercial-No Derivative Works 3.0
adamc@62 5 * Unported License.
adamc@62 6 * The license text is available at:
adamc@62 7 * http://creativecommons.org/licenses/by-nc-nd/3.0/
adamc@62 8 *)
adamc@62 9
adamc@62 10 (* begin hide *)
adamc@62 11 Require Import List.
adamc@62 12
adam@314 13 Require Import CpdtTactics.
adamc@62 14
adam@422 15 Definition bad : unit := tt.
adam@402 16
adamc@62 17 Set Implicit Arguments.
adamc@62 18 (* end hide *)
adamc@62 19
adamc@62 20
adamc@62 21 (** %\chapter{Infinite Data and Proofs}% *)
adamc@62 22
adam@395 23 (** In lazy functional programming languages like %\index{Haskell}%Haskell, infinite data structures are everywhere%~\cite{whyfp}%. Infinite lists and more exotic datatypes provide convenient abstractions for communication between parts of a program. Achieving similar convenience without infinite lazy structures would, in many cases, require acrobatic inversions of control flow.
adamc@62 24
adam@346 25 %\index{laziness}%Laziness is easy to implement in Haskell, where all the definitions in a program may be thought of as mutually recursive. In such an unconstrained setting, it is easy to implement an infinite loop when you really meant to build an infinite list, where any finite prefix of the list should be forceable in finite time. Haskell programmers learn how to avoid such slip-ups. In Coq, such a laissez-faire policy is not good enough.
adamc@62 26
adam@398 27 We spent some time in the last chapter discussing the %\index{Curry-Howard correspondence}%Curry-Howard isomorphism, where proofs are identified with functional programs. In such a setting, infinite loops, intended or otherwise, are disastrous. If Coq allowed the full breadth of definitions that Haskell did, we could code up an infinite loop and use it to prove any proposition vacuously. That is, the addition of general recursion would make CIC _inconsistent_. For an arbitrary proposition [P], we could write:
adamc@202 28 [[
adamc@202 29 Fixpoint bad (u : unit) : P := bad u.
adamc@205 30 ]]
adamc@205 31
adam@442 32 This would leave us with [bad tt] as a proof of [P].
adamc@62 33
adamc@62 34 There are also algorithmic considerations that make universal termination very desirable. We have seen how tactics like [reflexivity] compare terms up to equivalence under computational rules. Calls to recursive, pattern-matching functions are simplified automatically, with no need for explicit proof steps. It would be very hard to hold onto that kind of benefit if it became possible to write non-terminating programs; we would be running smack into the halting problem.
adamc@62 35
adam@422 36 One solution is to use types to contain the possibility of non-termination. For instance, we can create a "non-termination monad," inside which we must write all of our general-recursive programs; several such approaches are surveyed in Chapter 7. This is a heavyweight solution, and so we would like to avoid it whenever possible.
adamc@62 37
adam@402 38 Luckily, Coq has special support for a class of lazy data structures that happens to contain most examples found in Haskell. That mechanism,%\index{co-inductive types}% _co-inductive types_, is the subject of this chapter. *)
adamc@62 39
adamc@62 40
adamc@62 41 (** * Computing with Infinite Data *)
adamc@62 42
adam@398 43 (** Let us begin with the most basic type of infinite data, _streams_, or lazy lists.%\index{Vernacular commands!CoInductive}% *)
adamc@62 44
adamc@62 45 Section stream.
adam@351 46 Variable A : Type.
adamc@62 47
adam@351 48 CoInductive stream : Type :=
adamc@62 49 | Cons : A -> stream -> stream.
adamc@62 50 End stream.
adamc@62 51
adam@422 52 (* begin hide *)
adam@437 53 (* begin thide *)
adam@422 54 CoInductive evilStream := Nil.
adam@437 55 (* end thide *)
adam@422 56 (* end hide *)
adam@422 57
adamc@62 58 (** The definition is surprisingly simple. Starting from the definition of [list], we just need to change the keyword [Inductive] to [CoInductive]. We could have left a [Nil] constructor in our definition, but we will leave it out to force all of our streams to be infinite.
adamc@62 59
adam@402 60 How do we write down a stream constant? Obviously simple application of constructors is not good enough, since we could only denote finite objects that way. Rather, whereas recursive definitions were necessary to _use_ values of recursive inductive types effectively, here we find that we need%\index{co-recursive definitions}% _co-recursive definitions_ to _build_ values of co-inductive types effectively.
adamc@62 61
adam@346 62 We can define a stream consisting only of zeroes.%\index{Vernacular commands!CoFixpoint}% *)
adamc@62 63
adamc@62 64 CoFixpoint zeroes : stream nat := Cons 0 zeroes.
adamc@62 65
adam@346 66 (* EX: Define a stream that alternates between [true] and [false]. *)
adam@346 67 (* begin thide *)
adam@346 68
adamc@62 69 (** We can also define a stream that alternates between [true] and [false]. *)
adamc@62 70
adam@346 71 CoFixpoint trues_falses : stream bool := Cons true falses_trues
adam@346 72 with falses_trues : stream bool := Cons false trues_falses.
adam@346 73 (* end thide *)
adamc@62 74
adamc@62 75 (** Co-inductive values are fair game as arguments to recursive functions, and we can use that fact to write a function to take a finite approximation of a stream. *)
adamc@62 76
adam@348 77 (* EX: Define a function to calculate a finite approximation of a stream, to a particular length. *)
adam@346 78 (* begin thide *)
adam@346 79
adamc@211 80 Fixpoint approx A (s : stream A) (n : nat) : list A :=
adamc@62 81 match n with
adamc@62 82 | O => nil
adamc@62 83 | S n' =>
adamc@62 84 match s with
adamc@62 85 | Cons h t => h :: approx t n'
adamc@62 86 end
adamc@62 87 end.
adamc@62 88
adamc@62 89 Eval simpl in approx zeroes 10.
adamc@211 90 (** %\vspace{-.15in}% [[
adamc@62 91 = 0 :: 0 :: 0 :: 0 :: 0 :: 0 :: 0 :: 0 :: 0 :: 0 :: nil
adamc@62 92 : list nat
adam@302 93 ]]
adam@302 94 *)
adamc@211 95
adam@346 96 Eval simpl in approx trues_falses 10.
adamc@211 97 (** %\vspace{-.15in}% [[
adamc@62 98 = true
adamc@62 99 :: false
adamc@62 100 :: true
adamc@62 101 :: false
adamc@62 102 :: true :: false :: true :: false :: true :: false :: nil
adamc@62 103 : list bool
adam@346 104 ]]
adam@346 105 *)
adamc@62 106
adam@349 107 (* end thide *)
adamc@62 108
adam@402 109 (* begin hide *)
adam@437 110 (* begin thide *)
adam@402 111 Definition looper := 0.
adam@437 112 (* end thide *)
adam@402 113 (* end hide *)
adam@402 114
adam@398 115 (** So far, it looks like co-inductive types might be a magic bullet, allowing us to import all of the Haskeller's usual tricks. However, there are important restrictions that are dual to the restrictions on the use of inductive types. Fixpoints _consume_ values of inductive types, with restrictions on which _arguments_ may be passed in recursive calls. Dually, co-fixpoints _produce_ values of co-inductive types, with restrictions on what may be done with the _results_ of co-recursive calls.
adamc@62 116
adam@402 117 The restriction for co-inductive types shows up as the%\index{guardedness condition}% _guardedness condition_. First, consider this stream definition, which would be legal in Haskell.
adamc@62 118 [[
adamc@62 119 CoFixpoint looper : stream nat := looper.
adam@346 120 ]]
adamc@205 121
adam@346 122 <<
adamc@62 123 Error:
adamc@62 124 Recursive definition of looper is ill-formed.
adamc@62 125 In environment
adamc@62 126 looper : stream nat
adamc@62 127
adamc@62 128 unguarded recursive call in "looper"
adam@346 129 >>
adamc@205 130
adam@398 131 The rule we have run afoul of here is that _every co-recursive call must be guarded by a constructor_; that is, every co-recursive call must be a direct argument to a constructor of the co-inductive type we are generating. It is a good thing that this rule is enforced. If the definition of [looper] were accepted, our [approx] function would run forever when passed [looper], and we would have fallen into inconsistency.
adamc@62 132
adam@346 133 Some familiar functions are easy to write in co-recursive fashion. *)
adamc@62 134
adamc@62 135 Section map.
adam@351 136 Variables A B : Type.
adamc@62 137 Variable f : A -> B.
adamc@62 138
adamc@62 139 CoFixpoint map (s : stream A) : stream B :=
adamc@62 140 match s with
adamc@62 141 | Cons h t => Cons (f h) (map t)
adamc@62 142 end.
adamc@62 143 End map.
adamc@62 144
adam@402 145 (* begin hide *)
adam@437 146 (* begin thide *)
adam@402 147 Definition filter := 0.
adam@437 148 (* end thide *)
adam@402 149 (* end hide *)
adam@402 150
adam@402 151 (** This code is a literal copy of that for the list [map] function, with the [nil] case removed and [Fixpoint] changed to [CoFixpoint]. Many other standard functions on lazy data structures can be implemented just as easily. Some, like [filter], cannot be implemented. Since the predicate passed to [filter] may reject every element of the stream, we cannot satisfy the guardedness condition.
adamc@62 152
adam@398 153 The implications of the condition can be subtle. To illustrate how, we start off with another co-recursive function definition that _is_ legal. The function [interleave] takes two streams and produces a new stream that alternates between their elements. *)
adamc@62 154
adamc@62 155 Section interleave.
adam@351 156 Variable A : Type.
adamc@62 157
adamc@62 158 CoFixpoint interleave (s1 s2 : stream A) : stream A :=
adamc@62 159 match s1, s2 with
adamc@62 160 | Cons h1 t1, Cons h2 t2 => Cons h1 (Cons h2 (interleave t1 t2))
adamc@62 161 end.
adamc@62 162 End interleave.
adamc@62 163
adamc@62 164 (** Now say we want to write a weird stuttering version of [map] that repeats elements in a particular way, based on interleaving. *)
adamc@62 165
adamc@62 166 Section map'.
adam@351 167 Variables A B : Type.
adamc@62 168 Variable f : A -> B.
adamc@68 169 (* begin thide *)
adam@440 170 (** %\vspace{-.15in}%[[
adamc@62 171 CoFixpoint map' (s : stream A) : stream B :=
adamc@62 172 match s with
adam@346 173 | Cons h t => interleave (Cons (f h) (map' t)) (Cons (f h) (map' t))
adamc@68 174 end.
adamc@205 175 ]]
adam@440 176 %\vspace{-.15in}%We get another error message about an unguarded recursive call. *)
adamc@62 177
adam@346 178 End map'.
adam@346 179
adam@346 180 (** What is going wrong here? Imagine that, instead of [interleave], we had called some other, less well-behaved function on streams. Here is one simpler example demonstrating the essential pitfall. We start defining a standard function for taking the tail of a stream. Since streams are infinite, this operation is total. *)
adam@346 181
adam@346 182 Definition tl A (s : stream A) : stream A :=
adam@346 183 match s with
adam@346 184 | Cons _ s' => s'
adam@346 185 end.
adam@346 186
adam@346 187 (** Coq rejects the following definition that uses [tl].
adam@346 188 [[
adam@346 189 CoFixpoint bad : stream nat := tl (Cons 0 bad).
adam@346 190 ]]
adam@346 191
adam@422 192 Imagine that Coq had accepted our definition, and consider how we might evaluate [approx bad 1]. We would be trying to calculate the first element in the stream [bad]. However, it is not hard to see that the definition of [bad] "begs the question": unfolding the definition of [tl], we see that we essentially say "define [bad] to equal itself"! Of course such an equation admits no single well-defined solution, which does not fit well with the determinism of Gallina reduction.
adam@346 193
adam@346 194 Since Coq can be considered to check definitions after inlining and simplification of previously defined identifiers, the basic guardedness condition rules out our definition of [bad]. Such an inlining reduces [bad] to:
adam@346 195 [[
adam@346 196 CoFixpoint bad : stream nat := bad.
adam@346 197 ]]
adam@346 198 This is the same looping definition we rejected earlier. A similar inlining process reveals the way that Coq saw our failed definition of [map']:
adam@346 199 [[
adam@346 200 CoFixpoint map' (s : stream A) : stream B :=
adam@346 201 match s with
adam@346 202 | Cons h t => Cons (f h) (Cons (f h) (interleave (map' t) (map' t)))
adam@346 203 end.
adam@346 204 ]]
adam@346 205 Clearly in this case the [map'] calls are not immediate arguments to constructors, so we violate the guardedness condition. *)
adamc@68 206 (* end thide *)
adamc@211 207
adam@402 208 (** A more interesting question is why that condition is the right one. We can make an intuitive argument that the original [map'] definition is perfectly reasonable and denotes a well-understood transformation on streams, such that every output would behave properly with [approx]. The guardedness condition is an example of a syntactic check for%\index{productivity}% _productivity_ of co-recursive definitions. A productive definition can be thought of as one whose outputs can be forced in finite time to any finite approximation level, as with [approx]. If we replaced the guardedness condition with more involved checks, we might be able to detect and allow a broader range of productive definitions. However, mistakes in these checks could cause inconsistency, and programmers would need to understand the new, more complex checks. Coq's design strikes a balance between consistency and simplicity with its choice of guard condition, though we can imagine other worthwhile balances being struck, too. *)
adamc@62 209
adamc@63 210
adamc@63 211 (** * Infinite Proofs *)
adamc@63 212
adamc@63 213 (** Let us say we want to give two different definitions of a stream of all ones, and then we want to prove that they are equivalent. *)
adamc@63 214
adamc@63 215 CoFixpoint ones : stream nat := Cons 1 ones.
adamc@63 216 Definition ones' := map S zeroes.
adamc@63 217
adamc@63 218 (** The obvious statement of the equality is this: *)
adamc@63 219
adamc@63 220 Theorem ones_eq : ones = ones'.
adamc@63 221
adam@422 222 (* begin hide *)
adam@437 223 (* begin thide *)
adam@422 224 Definition foo := @eq.
adam@437 225 (* end thide *)
adam@422 226 (* end hide *)
adam@422 227
adamc@63 228 (** However, faced with the initial subgoal, it is not at all clear how this theorem can be proved. In fact, it is unprovable. The [eq] predicate that we use is fundamentally limited to equalities that can be demonstrated by finite, syntactic arguments. To prove this equivalence, we will need to introduce a new relation. *)
adamc@68 229 (* begin thide *)
adamc@211 230
adamc@63 231 Abort.
adamc@63 232
adam@398 233 (** Co-inductive datatypes make sense by analogy from Haskell. What we need now is a _co-inductive proposition_. That is, we want to define a proposition whose proofs may be infinite, subject to the guardedness condition. The idea of infinite proofs does not show up in usual mathematics, but it can be very useful (unsurprisingly) for reasoning about infinite data structures. Besides examples from Haskell, infinite data and proofs will also turn out to be useful for modelling inherently infinite mathematical objects, like program executions.
adamc@63 234
adam@346 235 We are ready for our first %\index{co-inductive predicates}%co-inductive predicate. *)
adamc@63 236
adamc@63 237 Section stream_eq.
adam@351 238 Variable A : Type.
adamc@63 239
adamc@63 240 CoInductive stream_eq : stream A -> stream A -> Prop :=
adamc@63 241 | Stream_eq : forall h t1 t2,
adamc@63 242 stream_eq t1 t2
adamc@63 243 -> stream_eq (Cons h t1) (Cons h t2).
adamc@63 244 End stream_eq.
adamc@63 245
adamc@63 246 (** We say that two streams are equal if and only if they have the same heads and their tails are equal. We use the normal finite-syntactic equality for the heads, and we refer to our new equality recursively for the tails.
adamc@63 247
adamc@63 248 We can try restating the theorem with [stream_eq]. *)
adamc@63 249
adamc@63 250 Theorem ones_eq : stream_eq ones ones'.
adamc@63 251 (** Coq does not support tactical co-inductive proofs as well as it supports tactical inductive proofs. The usual starting point is the [cofix] tactic, which asks to structure this proof as a co-fixpoint. *)
adamc@211 252
adamc@63 253 cofix.
adamc@63 254 (** [[
adamc@63 255 ones_eq : stream_eq ones ones'
adamc@63 256 ============================
adamc@63 257 stream_eq ones ones'
adamc@211 258
adamc@211 259 ]]
adamc@63 260
adamc@211 261 It looks like this proof might be easier than we expected! *)
adamc@63 262
adamc@63 263 assumption.
adam@422 264 (** <<
adamc@211 265 Proof completed.
adam@422 266 >>
adamc@63 267
adamc@211 268 Unfortunately, we are due for some disappointment in our victory lap.
adamc@211 269 [[
adamc@63 270 Qed.
adam@346 271 ]]
adamc@63 272
adam@346 273 <<
adamc@63 274 Error:
adamc@63 275 Recursive definition of ones_eq is ill-formed.
adamc@63 276
adamc@63 277 In environment
adamc@63 278 ones_eq : stream_eq ones ones'
adamc@63 279
adamc@205 280 unguarded recursive call in "ones_eq"
adam@346 281 >>
adamc@205 282
adamc@211 283 Via the Curry-Howard correspondence, the same guardedness condition applies to our co-inductive proofs as to our co-inductive data structures. We should be grateful that this proof is rejected, because, if it were not, the same proof structure could be used to prove any co-inductive theorem vacuously, by direct appeal to itself!
adamc@63 284
adam@347 285 Thinking about how Coq would generate a proof term from the proof script above, we see that the problem is that we are violating the guardedness condition. During our proofs, Coq can help us check whether we have yet gone wrong in this way. We can run the command [Guarded] in any context to see if it is possible to finish the proof in a way that will yield a properly guarded proof term.%\index{Vernacular commands!Guarded}%
adamc@63 286 [[
adamc@63 287 Guarded.
adamc@205 288 ]]
adamc@205 289
adam@398 290 Running [Guarded] here gives us the same error message that we got when we tried to run [Qed]. In larger proofs, [Guarded] can be helpful in detecting problems _before_ we think we are ready to run [Qed].
adamc@63 291
adam@281 292 We need to start the co-induction by applying [stream_eq]'s constructor. To do that, we need to know that both arguments to the predicate are [Cons]es. Informally, this is trivial, but [simpl] is not able to help us. *)
adamc@63 293
adamc@63 294 Undo.
adamc@63 295 simpl.
adamc@63 296 (** [[
adamc@63 297 ones_eq : stream_eq ones ones'
adamc@63 298 ============================
adamc@63 299 stream_eq ones ones'
adamc@211 300
adamc@211 301 ]]
adamc@63 302
adamc@211 303 It turns out that we are best served by proving an auxiliary lemma. *)
adamc@211 304
adamc@63 305 Abort.
adamc@63 306
adamc@63 307 (** First, we need to define a function that seems pointless on first glance. *)
adamc@63 308
adamc@63 309 Definition frob A (s : stream A) : stream A :=
adamc@63 310 match s with
adamc@63 311 | Cons h t => Cons h t
adamc@63 312 end.
adamc@63 313
adamc@63 314 (** Next, we need to prove a theorem that seems equally pointless. *)
adamc@63 315
adamc@63 316 Theorem frob_eq : forall A (s : stream A), s = frob s.
adamc@63 317 destruct s; reflexivity.
adamc@63 318 Qed.
adamc@63 319
adamc@63 320 (** But, miraculously, this theorem turns out to be just what we needed. *)
adamc@63 321
adamc@63 322 Theorem ones_eq : stream_eq ones ones'.
adamc@63 323 cofix.
adamc@63 324
adamc@63 325 (** We can use the theorem to rewrite the two streams. *)
adamc@211 326
adamc@63 327 rewrite (frob_eq ones).
adamc@63 328 rewrite (frob_eq ones').
adamc@63 329 (** [[
adamc@63 330 ones_eq : stream_eq ones ones'
adamc@63 331 ============================
adamc@63 332 stream_eq (frob ones) (frob ones')
adamc@211 333
adamc@211 334 ]]
adamc@63 335
adamc@211 336 Now [simpl] is able to reduce the streams. *)
adamc@63 337
adamc@63 338 simpl.
adamc@63 339 (** [[
adamc@63 340 ones_eq : stream_eq ones ones'
adamc@63 341 ============================
adamc@63 342 stream_eq (Cons 1 ones)
adamc@63 343 (Cons 1
adamc@63 344 ((cofix map (s : stream nat) : stream nat :=
adamc@63 345 match s with
adamc@63 346 | Cons h t => Cons (S h) (map t)
adamc@63 347 end) zeroes))
adamc@211 348
adamc@211 349 ]]
adamc@63 350
adam@422 351 Note the [cofix] notation for anonymous co-recursion, which is analogous to the [fix] notation we have already seen for recursion. Since we have exposed the [Cons] structure of each stream, we can apply the constructor of [stream_eq]. *)
adamc@63 352
adamc@63 353 constructor.
adamc@63 354 (** [[
adamc@63 355 ones_eq : stream_eq ones ones'
adamc@63 356 ============================
adamc@63 357 stream_eq ones
adamc@63 358 ((cofix map (s : stream nat) : stream nat :=
adamc@63 359 match s with
adamc@63 360 | Cons h t => Cons (S h) (map t)
adamc@63 361 end) zeroes)
adamc@211 362
adamc@211 363 ]]
adamc@63 364
adamc@211 365 Now, modulo unfolding of the definition of [map], we have matched our assumption. *)
adamc@211 366
adamc@63 367 assumption.
adamc@63 368 Qed.
adamc@63 369
adamc@63 370 (** Why did this silly-looking trick help? The answer has to do with the constraints placed on Coq's evaluation rules by the need for termination. The [cofix]-related restriction that foiled our first attempt at using [simpl] is dual to a restriction for [fix]. In particular, an application of an anonymous [fix] only reduces when the top-level structure of the recursive argument is known. Otherwise, we would be unfolding the recursive definition ad infinitum.
adamc@63 371
adam@398 372 Fixpoints only reduce when enough is known about the _definitions_ of their arguments. Dually, co-fixpoints only reduce when enough is known about _how their results will be used_. In particular, a [cofix] is only expanded when it is the discriminee of a [match]. Rewriting with our superficially silly lemma wrapped new [match]es around the two [cofix]es, triggering reduction.
adamc@63 373
adamc@63 374 If [cofix]es reduced haphazardly, it would be easy to run into infinite loops in evaluation, since we are, after all, building infinite objects.
adamc@63 375
adamc@63 376 One common source of difficulty with co-inductive proofs is bad interaction with standard Coq automation machinery. If we try to prove [ones_eq'] with automation, like we have in previous inductive proofs, we get an invalid proof. *)
adamc@63 377
adamc@63 378 Theorem ones_eq' : stream_eq ones ones'.
adamc@63 379 cofix; crush.
adam@440 380 (** %\vspace{-.25in}%[[
adamc@205 381 Guarded.
adam@302 382 ]]
adam@440 383 %\vspace{-.25in}%
adam@302 384 *)
adamc@63 385 Abort.
adam@346 386
adam@422 387 (** The standard [auto] machinery sees that our goal matches an assumption and so applies that assumption, even though this violates guardedness. One usually starts a proof like this by [destruct]ing some parameter and running a custom tactic to figure out the first proof rule to apply for each case. Alternatively, there are tricks that can be played with "hiding" the co-inductive hypothesis.
adam@346 388
adam@346 389 %\medskip%
adam@346 390
adam@402 391 Must we always be cautious with automation in proofs by co-induction? Induction seems to have dual versions of the same pitfalls inherent in it, and yet we avoid those pitfalls by encapsulating safe Curry-Howard recursion schemes inside named induction principles. It turns out that we can usually do the same with%\index{co-induction principles}% _co-induction principles_. Let us take that tack here, so that we can arrive at an [induction x; crush]-style proof for [ones_eq'].
adam@346 392
adam@398 393 An induction principle is parameterized over a predicate characterizing what we mean to prove, _as a function of the inductive fact that we already know_. Dually, a co-induction principle ought to be parameterized over a predicate characterizing what we mean to prove, _as a function of the arguments to the co-inductive predicate that we are trying to prove_.
adam@346 394
adam@346 395 To state a useful principle for [stream_eq], it will be useful first to define the stream head function. *)
adam@346 396
adam@346 397 Definition hd A (s : stream A) : A :=
adam@346 398 match s with
adam@346 399 | Cons x _ => x
adam@346 400 end.
adam@346 401
adam@346 402 (** Now we enter a section for the co-induction principle, based on %\index{Park's principle}%Park's principle as introduced in a tutorial by Gim%\'%enez%~\cite{IT}%. *)
adam@346 403
adam@346 404 Section stream_eq_coind.
adam@351 405 Variable A : Type.
adam@346 406 Variable R : stream A -> stream A -> Prop.
adam@392 407 (** This relation generalizes the theorem we want to prove, characterizing exactly which two arguments to [stream_eq] we want to consider. *)
adam@346 408
adam@346 409 Hypothesis Cons_case_hd : forall s1 s2, R s1 s2 -> hd s1 = hd s2.
adam@346 410 Hypothesis Cons_case_tl : forall s1 s2, R s1 s2 -> R (tl s1) (tl s2).
adam@422 411 (** Two hypotheses characterize what makes a good choice of [R]: it enforces equality of stream heads, and it is %``%#<i>#hereditary#</i>#%''% in the sense that an [R] stream pair passes on "[R]-ness" to its tails. An established technical term for such a relation is%\index{bisimulation}% _bisimulation_. *)
adam@346 412
adam@346 413 (** Now it is straightforward to prove the principle, which says that any stream pair in [R] is equal. The reader may wish to step through the proof script to see what is going on. *)
adam@392 414
adam@346 415 Theorem stream_eq_coind : forall s1 s2, R s1 s2 -> stream_eq s1 s2.
adam@346 416 cofix; destruct s1; destruct s2; intro.
adam@346 417 generalize (Cons_case_hd H); intro Heq; simpl in Heq; rewrite Heq.
adam@346 418 constructor.
adam@346 419 apply stream_eq_coind.
adam@346 420 apply (Cons_case_tl H).
adam@346 421 Qed.
adam@346 422 End stream_eq_coind.
adam@346 423
adam@346 424 (** To see why this proof is guarded, we can print it and verify that the one co-recursive call is an immediate argument to a constructor. *)
adam@392 425
adam@346 426 Print stream_eq_coind.
adam@346 427
adam@346 428 (** We omit the output and proceed to proving [ones_eq''] again. The only bit of ingenuity is in choosing [R], and in this case the most restrictive predicate works. *)
adam@346 429
adam@346 430 Theorem ones_eq'' : stream_eq ones ones'.
adam@346 431 apply (stream_eq_coind (fun s1 s2 => s1 = ones /\ s2 = ones')); crush.
adam@346 432 Qed.
adam@346 433
adam@346 434 (** Note that this proof achieves the proper reduction behavior via [hd] and [tl], rather than [frob] as in the last proof. All three functions pattern match on their arguments, catalyzing computation steps.
adam@346 435
adam@346 436 Compared to the inductive proofs that we are used to, it still seems unsatisfactory that we had to write out a choice of [R] in the last proof. An alternate is to capture a common pattern of co-recursion in a more specialized co-induction principle. For the current example, that pattern is: prove [stream_eq s1 s2] where [s1] and [s2] are defined as their own tails. *)
adam@346 437
adam@346 438 Section stream_eq_loop.
adam@351 439 Variable A : Type.
adam@346 440 Variables s1 s2 : stream A.
adam@346 441
adam@346 442 Hypothesis Cons_case_hd : hd s1 = hd s2.
adam@346 443 Hypothesis loop1 : tl s1 = s1.
adam@346 444 Hypothesis loop2 : tl s2 = s2.
adam@346 445
adam@346 446 (** The proof of the principle includes a choice of [R], so that we no longer need to make such choices thereafter. *)
adam@346 447
adam@346 448 Theorem stream_eq_loop : stream_eq s1 s2.
adam@346 449 apply (stream_eq_coind (fun s1' s2' => s1' = s1 /\ s2' = s2)); crush.
adam@346 450 Qed.
adam@346 451 End stream_eq_loop.
adam@346 452
adam@346 453 Theorem ones_eq''' : stream_eq ones ones'.
adam@346 454 apply stream_eq_loop; crush.
adam@346 455 Qed.
adamc@68 456 (* end thide *)
adamc@63 457
adam@435 458 (** Let us put [stream_eq_coind] through its paces a bit more, considering two different ways to compute infinite streams of all factorial values. First, we import the [fact] factorial function from the standard library. *)
adam@346 459
adam@346 460 Require Import Arith.
adam@346 461 Print fact.
adam@346 462 (** %\vspace{-.15in}%[[
adam@346 463 fact =
adam@346 464 fix fact (n : nat) : nat :=
adam@346 465 match n with
adam@346 466 | 0 => 1
adam@346 467 | S n0 => S n0 * fact n0
adam@346 468 end
adam@346 469 : nat -> nat
adam@346 470 ]]
adam@346 471 *)
adam@346 472
adam@346 473 (** The simplest way to compute the factorial stream involves calling [fact] afresh at each position. *)
adam@346 474
adam@346 475 CoFixpoint fact_slow' (n : nat) := Cons (fact n) (fact_slow' (S n)).
adam@346 476 Definition fact_slow := fact_slow' 1.
adam@346 477
adam@346 478 (** A more clever, optimized method maintains an accumulator of the previous factorial, so that each new entry can be computed with a single multiplication. *)
adam@346 479
adam@346 480 CoFixpoint fact_iter' (cur acc : nat) := Cons acc (fact_iter' (S cur) (acc * cur)).
adam@346 481 Definition fact_iter := fact_iter' 2 1.
adam@346 482
adam@346 483 (** We can verify that the streams are equal up to particular finite bounds. *)
adam@346 484
adam@346 485 Eval simpl in approx fact_iter 5.
adam@346 486 (** %\vspace{-.15in}%[[
adam@346 487 = 1 :: 2 :: 6 :: 24 :: 120 :: nil
adam@346 488 : list nat
adam@346 489 ]]
adam@346 490 *)
adam@346 491 Eval simpl in approx fact_slow 5.
adam@346 492 (** %\vspace{-.15in}%[[
adam@346 493 = 1 :: 2 :: 6 :: 24 :: 120 :: nil
adam@346 494 : list nat
adam@346 495 ]]
adam@346 496
adam@346 497 Now, to prove that the two versions are equivalent, it is helpful to prove (and add as a proof hint) a quick lemma about the computational behavior of [fact]. *)
adam@346 498
adam@346 499 (* begin thide *)
adam@346 500 Lemma fact_def : forall x n,
adam@346 501 fact_iter' x (fact n * S n) = fact_iter' x (fact (S n)).
adam@346 502 simpl; intros; f_equal; ring.
adam@346 503 Qed.
adam@346 504
adam@346 505 Hint Resolve fact_def.
adam@346 506
adam@346 507 (** With the hint added, it is easy to prove an auxiliary lemma relating [fact_iter'] and [fact_slow']. The key bit of ingenuity is introduction of an existential quantifier for the shared parameter [n]. *)
adam@346 508
adam@346 509 Lemma fact_eq' : forall n, stream_eq (fact_iter' (S n) (fact n)) (fact_slow' n).
adam@346 510 intro; apply (stream_eq_coind (fun s1 s2 => exists n, s1 = fact_iter' (S n) (fact n)
adam@346 511 /\ s2 = fact_slow' n)); crush; eauto.
adam@346 512 Qed.
adam@346 513
adam@346 514 (** The final theorem is a direct corollary of [fact_eq']. *)
adam@346 515
adam@346 516 Theorem fact_eq : stream_eq fact_iter fact_slow.
adam@346 517 apply fact_eq'.
adam@346 518 Qed.
adam@346 519
adam@346 520 (** As in the case of [ones_eq'], we may be unsatisfied that we needed to write down a choice of [R] that seems to duplicate information already present in a lemma statement. We can facilitate a simpler proof by defining a co-induction principle specialized to goals that begin with single universal quantifiers, and the strategy can be extended in a straightforward way to principles for other counts of quantifiers. (Our [stream_eq_loop] principle is effectively the instantiation of this technique to zero quantifiers.) *)
adam@346 521
adam@346 522 Section stream_eq_onequant.
adam@351 523 Variables A B : Type.
adam@346 524 (** We have the types [A], the domain of the one quantifier; and [B], the type of data found in the streams. *)
adam@346 525
adam@346 526 Variables f g : A -> stream B.
adam@346 527 (** The two streams we compare must be of the forms [f x] and [g x], for some shared [x]. Note that this falls out naturally when [x] is a shared universally quantified variable in a lemma statement. *)
adam@346 528
adam@346 529 Hypothesis Cons_case_hd : forall x, hd (f x) = hd (g x).
adam@346 530 Hypothesis Cons_case_tl : forall x, exists y, tl (f x) = f y /\ tl (g x) = g y.
adam@346 531 (** These conditions are inspired by the bisimulation requirements, with a more general version of the [R] choice we made for [fact_eq'] inlined into the hypotheses of [stream_eq_coind]. *)
adam@346 532
adam@346 533 Theorem stream_eq_onequant : forall x, stream_eq (f x) (g x).
adam@346 534 intro; apply (stream_eq_coind (fun s1 s2 => exists x, s1 = f x /\ s2 = g x)); crush; eauto.
adam@346 535 Qed.
adam@346 536 End stream_eq_onequant.
adam@346 537
adam@346 538 Lemma fact_eq'' : forall n, stream_eq (fact_iter' (S n) (fact n)) (fact_slow' n).
adam@346 539 apply stream_eq_onequant; crush; eauto.
adam@346 540 Qed.
adam@346 541
adam@346 542 (** We have arrived at one of our customary automated proofs, thanks to the new principle. *)
adam@346 543 (* end thide *)
adamc@64 544
adamc@64 545
adamc@64 546 (** * Simple Modeling of Non-Terminating Programs *)
adamc@64 547
adam@402 548 (** We close the chapter with a quick motivating example for more complex uses of co-inductive types. We will define a co-inductive semantics for a simple imperative programming language and use that semantics to prove the correctness of a trivial optimization that removes spurious additions by 0. We follow the technique of%\index{co-inductive big-step operational semantics}% _co-inductive big-step operational semantics_ %\cite{BigStep}%.
adamc@64 549
adam@347 550 We define a suggestive synonym for [nat], as we will consider programs with infinitely many variables, represented as [nat]s. *)
adamc@211 551
adam@347 552 Definition var := nat.
adamc@64 553
adam@436 554 (** We define a type [vars] of maps from variables to values. To define a function [set] for setting a variable's value in a map, we use the standard library function [beq_nat] for comparing natural numbers. *)
adamc@64 555
adam@347 556 Definition vars := var -> nat.
adam@347 557 Definition set (vs : vars) (v : var) (n : nat) : vars :=
adam@347 558 fun v' => if beq_nat v v' then n else vs v'.
adamc@67 559
adam@347 560 (** We define a simple arithmetic expression language with variables, and we give it a semantics via an interpreter. *)
adamc@67 561
adam@347 562 Inductive exp : Set :=
adam@347 563 | Const : nat -> exp
adam@347 564 | Var : var -> exp
adam@347 565 | Plus : exp -> exp -> exp.
adamc@64 566
adam@347 567 Fixpoint evalExp (vs : vars) (e : exp) : nat :=
adam@347 568 match e with
adam@347 569 | Const n => n
adam@347 570 | Var v => vs v
adam@347 571 | Plus e1 e2 => evalExp vs e1 + evalExp vs e2
adam@347 572 end.
adamc@64 573
adam@422 574 (** Finally, we define a language of commands. It includes variable assignment, sequencing, and a <<while>> form that repeats as long as its test expression evaluates to a nonzero value. *)
adamc@64 575
adam@347 576 Inductive cmd : Set :=
adam@347 577 | Assign : var -> exp -> cmd
adam@347 578 | Seq : cmd -> cmd -> cmd
adam@347 579 | While : exp -> cmd -> cmd.
adamc@64 580
adam@398 581 (** We could define an inductive relation to characterize the results of command evaluation. However, such a relation would not capture _nonterminating_ executions. With a co-inductive relation, we can capture both cases. The parameters of the relation are an initial state, a command, and a final state. A program that does not terminate in a particular initial state is related to _any_ final state. *)
adamc@67 582
adam@347 583 CoInductive evalCmd : vars -> cmd -> vars -> Prop :=
adam@347 584 | EvalAssign : forall vs v e, evalCmd vs (Assign v e) (set vs v (evalExp vs e))
adam@347 585 | EvalSeq : forall vs1 vs2 vs3 c1 c2, evalCmd vs1 c1 vs2
adam@347 586 -> evalCmd vs2 c2 vs3
adam@347 587 -> evalCmd vs1 (Seq c1 c2) vs3
adam@347 588 | EvalWhileFalse : forall vs e c, evalExp vs e = 0
adam@347 589 -> evalCmd vs (While e c) vs
adam@347 590 | EvalWhileTrue : forall vs1 vs2 vs3 e c, evalExp vs1 e <> 0
adam@347 591 -> evalCmd vs1 c vs2
adam@347 592 -> evalCmd vs2 (While e c) vs3
adam@347 593 -> evalCmd vs1 (While e c) vs3.
adam@347 594
adam@347 595 (** Having learned our lesson in the last section, before proceeding, we build a co-induction principle for [evalCmd]. *)
adam@347 596
adam@347 597 Section evalCmd_coind.
adam@347 598 Variable R : vars -> cmd -> vars -> Prop.
adam@347 599
adam@347 600 Hypothesis AssignCase : forall vs1 vs2 v e, R vs1 (Assign v e) vs2
adam@347 601 -> vs2 = set vs1 v (evalExp vs1 e).
adam@347 602
adam@347 603 Hypothesis SeqCase : forall vs1 vs3 c1 c2, R vs1 (Seq c1 c2) vs3
adam@347 604 -> exists vs2, R vs1 c1 vs2 /\ R vs2 c2 vs3.
adam@347 605
adam@347 606 Hypothesis WhileCase : forall vs1 vs3 e c, R vs1 (While e c) vs3
adam@347 607 -> (evalExp vs1 e = 0 /\ vs3 = vs1)
adam@347 608 \/ exists vs2, evalExp vs1 e <> 0 /\ R vs1 c vs2 /\ R vs2 (While e c) vs3.
adam@347 609
adam@402 610 (** The proof is routine. We make use of a form of %\index{tactics!destruct}%[destruct] that takes an%\index{intro pattern}% _intro pattern_ in an [as] clause. These patterns control how deeply we break apart the components of an inductive value, and we refer the reader to the Coq manual for more details. *)
adam@347 611
adam@347 612 Theorem evalCmd_coind : forall vs1 c vs2, R vs1 c vs2 -> evalCmd vs1 c vs2.
adam@347 613 cofix; intros; destruct c.
adam@347 614 rewrite (AssignCase H); constructor.
adam@347 615 destruct (SeqCase H) as [? [? ?]]; econstructor; eauto.
adam@347 616 destruct (WhileCase H) as [[? ?] | [? [? [? ?]]]]; subst;
adam@347 617 [ econstructor | econstructor 4 ]; eauto.
adam@347 618 Qed.
adam@347 619 End evalCmd_coind.
adam@347 620
adam@347 621 (** Now that we have a co-induction principle, we should use it to prove something! Our example is a trivial program optimizer that finds places to replace [0 + e] with [e]. *)
adam@347 622
adam@347 623 Fixpoint optExp (e : exp) : exp :=
adam@347 624 match e with
adam@347 625 | Plus (Const 0) e => optExp e
adam@347 626 | Plus e1 e2 => Plus (optExp e1) (optExp e2)
adam@347 627 | _ => e
adam@347 628 end.
adam@347 629
adam@347 630 Fixpoint optCmd (c : cmd) : cmd :=
adam@347 631 match c with
adam@347 632 | Assign v e => Assign v (optExp e)
adam@347 633 | Seq c1 c2 => Seq (optCmd c1) (optCmd c2)
adam@347 634 | While e c => While (optExp e) (optCmd c)
adam@347 635 end.
adam@347 636
adam@347 637 (** Before proving correctness of [optCmd], we prove a lemma about [optExp]. This is where we have to do the most work, choosing pattern match opportunities automatically. *)
adam@347 638
adam@347 639 (* begin thide *)
adam@347 640 Lemma optExp_correct : forall vs e, evalExp vs (optExp e) = evalExp vs e.
adam@347 641 induction e; crush;
adam@347 642 repeat (match goal with
adam@347 643 | [ |- context[match ?E with Const _ => _ | Var _ => _
adam@347 644 | Plus _ _ => _ end] ] => destruct E
adam@347 645 | [ |- context[match ?E with O => _ | S _ => _ end] ] => destruct E
adam@347 646 end; crush).
adamc@64 647 Qed.
adamc@64 648
adam@375 649 Hint Rewrite optExp_correct .
adamc@64 650
adam@384 651 (** The final theorem is easy to establish, using our co-induction principle and a bit of Ltac smarts that we leave unexplained for now. Curious readers can consult the Coq manual, or wait for the later chapters of this book about proof automation. At a high level, we show inclusions between behaviors, going in both directions between original and optimized programs. *)
adamc@64 652
adam@384 653 Ltac finisher := match goal with
adam@384 654 | [ H : evalCmd _ _ _ |- _ ] => ((inversion H; [])
adam@384 655 || (inversion H; [|])); subst
adam@384 656 end; crush; eauto 10.
adam@384 657
adam@384 658 Lemma optCmd_correct1 : forall vs1 c vs2, evalCmd vs1 c vs2
adam@347 659 -> evalCmd vs1 (optCmd c) vs2.
adam@347 660 intros; apply (evalCmd_coind (fun vs1 c' vs2 => exists c, evalCmd vs1 c vs2
adam@347 661 /\ c' = optCmd c)); eauto; crush;
adam@347 662 match goal with
adam@347 663 | [ H : _ = optCmd ?E |- _ ] => destruct E; simpl in *; discriminate
adam@347 664 || injection H; intros; subst
adam@384 665 end; finisher.
adam@384 666 Qed.
adam@384 667
adam@384 668 Lemma optCmd_correct2 : forall vs1 c vs2, evalCmd vs1 (optCmd c) vs2
adam@384 669 -> evalCmd vs1 c vs2.
adam@384 670 intros; apply (evalCmd_coind (fun vs1 c vs2 => evalCmd vs1 (optCmd c) vs2));
adam@384 671 crush; finisher.
adam@384 672 Qed.
adam@384 673
adam@384 674 Theorem optCmd_correct : forall vs1 c vs2, evalCmd vs1 (optCmd c) vs2
adam@384 675 <-> evalCmd vs1 c vs2.
adam@384 676 intuition; apply optCmd_correct1 || apply optCmd_correct2; assumption.
adam@347 677 Qed.
adam@347 678 (* end thide *)
adamc@64 679
adam@422 680 (** In this form, the theorem tells us that the optimizer preserves observable behavior of both terminating and nonterminating programs, but we did not have to do more work than for the case of terminating programs alone. We merely took the natural inductive definition for terminating executions, made it co-inductive, and applied the appropriate co-induction principle. Curious readers might experiment with adding command constructs like <<if>>; the same proof script should continue working, after the co-induction principle is extended to the new evaluation rules. *)