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Chapter read-through
author Adam Chlipala <adamc@hcoop.net>
date Sun, 28 Sep 2008 13:50:21 -0400
parents 31622083ad5d
children 1946586b9da7
rev   line source
adamc@45 1 (* Copyright (c) 2008, Adam Chlipala
adamc@45 2 *
adamc@45 3 * This work is licensed under a
adamc@45 4 * Creative Commons Attribution-Noncommercial-No Derivative Works 3.0
adamc@45 5 * Unported License.
adamc@45 6 * The license text is available at:
adamc@45 7 * http://creativecommons.org/licenses/by-nc-nd/3.0/
adamc@45 8 *)
adamc@45 9
adamc@45 10 (* begin hide *)
adamc@45 11 Require Import List.
adamc@45 12
adamc@45 13 Require Import Tactics.
adamc@45 14
adamc@45 15 Set Implicit Arguments.
adamc@45 16 (* end hide *)
adamc@45 17
adamc@45 18
adamc@45 19 (** %\chapter{Inductive Predicates}% *)
adamc@45 20
adamc@45 21 (** The so-called "Curry-Howard Correspondence" states a formal connection between functional programs and mathematical proofs. In the last chapter, we snuck in a first introduction to this subject in Coq. Witness the close similarity between the types [unit] and [True] from the standard library: *)
adamc@45 22
adamc@45 23 Print unit.
adamc@45 24 (** [[
adamc@45 25
adamc@45 26 Inductive unit : Set := tt : unit
adamc@45 27 ]] *)
adamc@45 28
adamc@45 29 Print True.
adamc@45 30 (** [[
adamc@45 31
adamc@45 32 Inductive True : Prop := I : True
adamc@45 33 ]] *)
adamc@45 34
adamc@45 35 (** Recall that [unit] is the type with only one value, and [True] is the proposition that always holds. Despite this superficial difference between the two concepts, in both cases we can use the same inductive definition mechanism. The connection goes further than this. We see that we arrive at the definition of [True] by replacing [unit] by [True], [tt] by [I], and [Set] by [Prop]. The first two of these differences are superficial changes of names, while the third difference is the crucial one for separating programs from proofs. A term [T] of type [Set] is a type of programs, and a term of type [T] is a program. A term [T] of type [Prop] is a logical proposition, and its proofs are of type [T].
adamc@45 36
adamc@45 37 [unit] has one value, [tt]. [True] has one proof, [I]. Why distinguish between these two types? Many people who have read about Curry-Howard in an abstract context and not put it to use in proof engineering answer that the two types in fact %\textit{%#<i>#should not#</i>#%}% be distinguished. There is a certain aesthetic appeal to this point of view, but I want to argue that it is best to treat Curry-Howard very loosely in practical proving. There are Coq-specific reasons for preferring the distinction, involving efficient compilation and avoidance of paradoxes in the presence of classical math, but I will argue that there is a more general principle that should lead us to avoid conflating programming and proving.
adamc@45 38
adamc@45 39 The essence of the argument is roughly this: to an engineer, not all functions of type [A -> B] are created equal, but all proofs of a proposition [P -> Q] are. This idea is known as %\textit{%#<i>#proof irrelevance#</i>#%}%, and its formalizations in logics prevent us from distinguishing between alternate proofs of the same proposition. Proof irrelevance is compatible with, but not derivable in, Gallina. Apart from this theoretical concern, I will argue that it is most effective to do engineering with Coq by employing different techniques for programs versus proofs. Most of this book is organized around that distinction, describing how to program, by applying standard functional programming techniques in the presence of dependent types; and how to prove, by writing custom Ltac decision procedures.
adamc@45 40
adamc@45 41 With that perspective in mind, this chapter is sort of a mirror image of the last chapter, introducing how to define predicates with inductive definitions. We will point out similarities in places, but much of the effective Coq user's bag of tricks is disjoint for predicates versus "datatypes." This chapter is also a covert introduction to dependent types, which are the foundation on which interesting inductive predicates are built, though we will rely on tactics to build dependently-typed proof terms for us for now. A future chapter introduces more manual application of dependent types. *)
adamc@45 42
adamc@45 43
adamc@48 44 (** * Propositional Logic *)
adamc@45 45
adamc@45 46 (** Let us begin with a brief tour through the definitions of the connectives for propositional logic. We will work within a Coq section that provides us with a set of propositional variables. In Coq parlance, these are just terms of type [Prop.] *)
adamc@45 47
adamc@45 48 Section Propositional.
adamc@46 49 Variables P Q R : Prop.
adamc@45 50
adamc@45 51 (** In Coq, the most basic propositional connective is implication, written [->], which we have already used in almost every proof. Rather than being defined inductively, implication is built into Coq as the function type constructor.
adamc@45 52
adamc@45 53 We have also already seen the definition of [True]. For a demonstration of a lower-level way of establishing proofs of inductive predicates, we turn to this trivial theorem. *)
adamc@45 54
adamc@45 55 Theorem obvious : True.
adamc@55 56 (* begin thide *)
adamc@45 57 apply I.
adamc@55 58 (* end thide *)
adamc@45 59 Qed.
adamc@45 60
adamc@45 61 (** We may always use the [apply] tactic to take a proof step based on applying a particular constructor of the inductive predicate that we are trying to establish. Sometimes there is only one constructor that could possibly apply, in which case a shortcut is available: *)
adamc@45 62
adamc@55 63 (* begin thide *)
adamc@45 64 Theorem obvious' : True.
adamc@45 65 constructor.
adamc@45 66 Qed.
adamc@45 67
adamc@55 68 (* end thide *)
adamc@55 69
adamc@45 70 (** There is also a predicate [False], which is the Curry-Howard mirror image of [Empty_set] from the last chapter. *)
adamc@45 71
adamc@45 72 Print False.
adamc@45 73 (** [[
adamc@45 74
adamc@45 75 Inductive False : Prop :=
adamc@45 76 ]] *)
adamc@45 77
adamc@45 78 (** We can conclude anything from [False], doing case analysis on a proof of [False] in the same way we might do case analysis on, say, a natural number. Since there are no cases to consider, any such case analysis succeeds immediately in proving the goal. *)
adamc@45 79
adamc@45 80 Theorem False_imp : False -> 2 + 2 = 5.
adamc@55 81 (* begin thide *)
adamc@45 82 destruct 1.
adamc@55 83 (* end thide *)
adamc@45 84 Qed.
adamc@45 85
adamc@45 86 (** In a consistent context, we can never build a proof of [False]. In inconsistent contexts that appear in the courses of proofs, it is usually easiest to proceed by demonstrating that inconsistency with an explicit proof of [False]. *)
adamc@45 87
adamc@45 88 Theorem arith_neq : 2 + 2 = 5 -> 9 + 9 = 835.
adamc@55 89 (* begin thide *)
adamc@45 90 intro.
adamc@45 91
adamc@45 92 (** At this point, we have an inconsistent hypothesis [2 + 2 = 5], so the specific conclusion is not important. We use the [elimtype] tactic to state a proposition, telling Coq that we wish to construct a proof of the new proposition and then prove the original goal by case analysis on the structure of the new auxiliary proof. Since [False] has no constructors, [elimtype False] simply leaves us with the obligation to prove [False]. *)
adamc@45 93
adamc@45 94 elimtype False.
adamc@45 95 (** [[
adamc@45 96
adamc@45 97 H : 2 + 2 = 5
adamc@45 98 ============================
adamc@45 99 False
adamc@45 100 ]] *)
adamc@45 101
adamc@45 102 (** For now, we will leave the details of this proof about arithmetic to [crush]. *)
adamc@45 103
adamc@45 104 crush.
adamc@55 105 (* end thide *)
adamc@45 106 Qed.
adamc@45 107
adamc@45 108 (** A related notion to [False] is logical negation. *)
adamc@45 109
adamc@45 110 Print not.
adamc@45 111 (** [[
adamc@45 112
adamc@45 113 not = fun A : Prop => A -> False
adamc@45 114 : Prop -> Prop
adamc@45 115 ]] *)
adamc@45 116
adamc@45 117 (** We see that [not] is just shorthand for implication of [False]. We can use that fact explicitly in proofs. The syntax [~P] expands to [not P]. *)
adamc@45 118
adamc@45 119 Theorem arith_neq' : ~ (2 + 2 = 5).
adamc@55 120 (* begin thide *)
adamc@45 121 unfold not.
adamc@45 122
adamc@45 123 (** [[
adamc@45 124
adamc@45 125 ============================
adamc@45 126 2 + 2 = 5 -> False
adamc@45 127 ]] *)
adamc@45 128
adamc@45 129 crush.
adamc@55 130 (* end thide *)
adamc@45 131 Qed.
adamc@45 132
adamc@45 133 (** We also have conjunction, which we introduced in the last chapter. *)
adamc@45 134
adamc@45 135 Print and.
adamc@45 136 (** [[
adamc@45 137
adamc@45 138 Inductive and (A : Prop) (B : Prop) : Prop := conj : A -> B -> A /\ B
adamc@45 139 ]] *)
adamc@45 140
adamc@45 141 (** The interested reader can check that [and] has a Curry-Howard doppleganger called [prod], the type of pairs. However, it is generally most convenient to reason about conjunction using tactics. An explicit proof of commutativity of [and] illustrates the usual suspects for such tasks. [/\] is an infix shorthand for [and]. *)
adamc@45 142
adamc@45 143 Theorem and_comm : P /\ Q -> Q /\ P.
adamc@55 144 (* begin thide *)
adamc@45 145 (** We start by case analysis on the proof of [P /\ Q]. *)
adamc@45 146
adamc@45 147 destruct 1.
adamc@45 148 (** [[
adamc@45 149
adamc@45 150 H : P
adamc@45 151 H0 : Q
adamc@45 152 ============================
adamc@45 153 Q /\ P
adamc@45 154 ]] *)
adamc@45 155
adamc@45 156 (** Every proof of a conjunction provides proofs for both conjuncts, so we get a single subgoal reflecting that. We can proceed by splitting this subgoal into a case for each conjunct of [Q /\ P]. *)
adamc@45 157
adamc@45 158 split.
adamc@45 159 (** [[
adamc@45 160 2 subgoals
adamc@45 161
adamc@45 162 H : P
adamc@45 163 H0 : Q
adamc@45 164 ============================
adamc@45 165 Q
adamc@45 166
adamc@45 167 subgoal 2 is:
adamc@45 168 P
adamc@45 169 ]] *)
adamc@45 170
adamc@45 171 (** In each case, the conclusion is among our hypotheses, so the [assumption] tactic finishes the process. *)
adamc@45 172
adamc@45 173 assumption.
adamc@45 174 assumption.
adamc@55 175 (* end thide *)
adamc@45 176 Qed.
adamc@45 177
adamc@45 178 (** Coq disjunction is called [or] and abbreviated with the infix operator [\/]. *)
adamc@45 179
adamc@45 180 Print or.
adamc@45 181 (** [[
adamc@45 182
adamc@45 183 Inductive or (A : Prop) (B : Prop) : Prop :=
adamc@45 184 or_introl : A -> A \/ B | or_intror : B -> A \/ B
adamc@45 185 ]] *)
adamc@45 186
adamc@45 187 (** We see that there are two ways to prove a disjunction: prove the first disjunct or prove the second. The Curry-Howard analogue of this is the Coq [sum] type. We can demonstrate the main tactics here with another proof of commutativity. *)
adamc@45 188
adamc@45 189 Theorem or_comm : P \/ Q -> Q \/ P.
adamc@55 190
adamc@55 191 (* begin thide *)
adamc@45 192 (** As in the proof for [and], we begin with case analysis, though this time we are met by two cases instead of one. *)
adamc@45 193 destruct 1.
adamc@45 194 (** [[
adamc@45 195
adamc@45 196 2 subgoals
adamc@45 197
adamc@45 198 H : P
adamc@45 199 ============================
adamc@45 200 Q \/ P
adamc@45 201
adamc@45 202 subgoal 2 is:
adamc@45 203 Q \/ P
adamc@45 204 ]] *)
adamc@45 205
adamc@45 206 (** We can see that, in the first subgoal, we want to prove the disjunction by proving its second disjunct. The [right] tactic telegraphs this intent. *)
adamc@45 207
adamc@45 208 right; assumption.
adamc@45 209
adamc@45 210 (** The second subgoal has a symmetric proof.
adamc@45 211
adamc@45 212 [[
adamc@45 213
adamc@45 214 1 subgoal
adamc@45 215
adamc@45 216 H : Q
adamc@45 217 ============================
adamc@45 218 Q \/ P
adamc@45 219 ]] *)
adamc@45 220
adamc@45 221 left; assumption.
adamc@55 222 (* end thide *)
adamc@45 223 Qed.
adamc@45 224
adamc@46 225
adamc@46 226 (* begin hide *)
adamc@46 227 (* In-class exercises *)
adamc@46 228
adamc@46 229 Theorem contra : P -> ~P -> R.
adamc@52 230 (* begin thide *)
adamc@52 231 unfold not.
adamc@52 232 intros.
adamc@52 233 elimtype False.
adamc@52 234 apply H0.
adamc@52 235 assumption.
adamc@52 236 (* end thide *)
adamc@46 237 Admitted.
adamc@46 238
adamc@46 239 Theorem and_assoc : (P /\ Q) /\ R -> P /\ (Q /\ R).
adamc@52 240 (* begin thide *)
adamc@52 241 intros.
adamc@52 242 destruct H.
adamc@52 243 destruct H.
adamc@52 244 split.
adamc@52 245 assumption.
adamc@52 246 split.
adamc@52 247 assumption.
adamc@52 248 assumption.
adamc@52 249 (* end thide *)
adamc@46 250 Admitted.
adamc@46 251
adamc@46 252 Theorem or_assoc : (P \/ Q) \/ R -> P \/ (Q \/ R).
adamc@52 253 (* begin thide *)
adamc@52 254 intros.
adamc@52 255 destruct H.
adamc@52 256 destruct H.
adamc@52 257 left.
adamc@52 258 assumption.
adamc@52 259 right.
adamc@52 260 left.
adamc@52 261 assumption.
adamc@52 262 right.
adamc@52 263 right.
adamc@52 264 assumption.
adamc@52 265 (* end thide *)
adamc@46 266 Admitted.
adamc@46 267
adamc@46 268 (* end hide *)
adamc@46 269
adamc@46 270
adamc@46 271 (** It would be a shame to have to plod manually through all proofs about propositional logic. Luckily, there is no need. One of the most basic Coq automation tactics is [tauto], which is a complete decision procedure for constructive propositional logic. (More on what "constructive" means in the next section.) We can use [tauto] to dispatch all of the purely propositional theorems we have proved so far. *)
adamc@46 272
adamc@46 273 Theorem or_comm' : P \/ Q -> Q \/ P.
adamc@55 274 (* begin thide *)
adamc@46 275 tauto.
adamc@55 276 (* end thide *)
adamc@46 277 Qed.
adamc@46 278
adamc@46 279 (** Sometimes propositional reasoning forms important plumbing for the proof of a theorem, but we still need to apply some other smarts about, say, arithmetic. [intuition] is a generalization of [tauto] that proves everything it can using propositional reasoning. When some goals remain, it uses propositional laws to simplify them as far as possible. Consider this example, which uses the list concatenation operator [++] from the standard library. *)
adamc@46 280
adamc@46 281 Theorem arith_comm : forall ls1 ls2 : list nat,
adamc@46 282 length ls1 = length ls2 \/ length ls1 + length ls2 = 6
adamc@46 283 -> length (ls1 ++ ls2) = 6 \/ length ls1 = length ls2.
adamc@55 284 (* begin thide *)
adamc@46 285 intuition.
adamc@46 286
adamc@46 287 (** A lot of the proof structure has been generated for us by [intuition], but the final proof depends on a fact about lists. The remaining subgoal hints at what cleverness we need to inject. *)
adamc@46 288
adamc@46 289 (** [[
adamc@46 290
adamc@46 291 ls1 : list nat
adamc@46 292 ls2 : list nat
adamc@46 293 H0 : length ls1 + length ls2 = 6
adamc@46 294 ============================
adamc@46 295 length (ls1 ++ ls2) = 6 \/ length ls1 = length ls2
adamc@46 296 ]] *)
adamc@46 297
adamc@46 298 (** We can see that we need a theorem about lengths of concatenated lists, which we proved last chapter and is also in the standard library. *)
adamc@46 299
adamc@46 300 rewrite app_length.
adamc@46 301 (** [[
adamc@46 302
adamc@46 303 ls1 : list nat
adamc@46 304 ls2 : list nat
adamc@46 305 H0 : length ls1 + length ls2 = 6
adamc@46 306 ============================
adamc@46 307 length ls1 + length ls2 = 6 \/ length ls1 = length ls2
adamc@46 308 ]] *)
adamc@46 309
adamc@46 310 (** Now the subgoal follows by purely propositional reasoning. That is, we could replace [length ls1 + length ls2 = 6] with [P] and [length ls1 = length ls2] with [Q] and arrive at a tautology of propositional logic. *)
adamc@46 311
adamc@46 312 tauto.
adamc@55 313 (* end thide *)
adamc@46 314 Qed.
adamc@46 315
adamc@46 316 (** [intuition] is one of the main bits of glue in the implementation of [crush], so, with a little help, we can get a short automated proof of the theorem. *)
adamc@46 317
adamc@55 318 (* begin thide *)
adamc@46 319 Theorem arith_comm' : forall ls1 ls2 : list nat,
adamc@46 320 length ls1 = length ls2 \/ length ls1 + length ls2 = 6
adamc@46 321 -> length (ls1 ++ ls2) = 6 \/ length ls1 = length ls2.
adamc@46 322 Hint Rewrite app_length : cpdt.
adamc@46 323
adamc@46 324 crush.
adamc@46 325 Qed.
adamc@55 326 (* end thide *)
adamc@46 327
adamc@45 328 End Propositional.
adamc@45 329
adamc@46 330
adamc@47 331 (** * What Does It Mean to Be Constructive? *)
adamc@46 332
adamc@47 333 (** One potential point of confusion in the presentation so far is the distinction between [bool] and [Prop]. [bool] is a datatype whose two values are [true] and [false], while [Prop] is a more primitive type that includes among its members [True] and [False]. Why not collapse these two concepts into one, and why must there be more than two states of mathematical truth?
adamc@46 334
adamc@47 335 The answer comes from the fact that Coq implements %\textit{%#<i>#constructive#</i>#%}% or %\textit{%#<i>#intuitionistic#</i>#%}% logic, in contrast to the %\textit{%#<i>#classical#</i>#%}% logic that you may be more familiar with. In constructive logic, classical tautologies like [~ ~P -> P] and [P \/ ~P] do not always hold. In general, we can only prove these tautologies when [P] is %\textit{%#<i>#decidable#</i>#%}%, in the sense of computability theory. The Curry-Howard encoding that Coq uses for [or] allows us to extract either a proof of [P] or a proof of [~P] from any proof of [P \/ ~P]. Since our proofs are just functional programs which we can run, this would give us a decision procedure for the halting problem, where the instantiations of [P] would be formulas like "this particular Turing machine halts."
adamc@47 336
adamc@47 337 Hence the distinction between [bool] and [Prop]. Programs of type [bool] are computational by construction; we can always run them to determine their results. Many [Prop]s are undecidable, and so we can write more expressive formulas with [Prop]s than with [bool]s, but the inevitable consequence is that we cannot simply "run a [Prop] to determine its truth."
adamc@47 338
adamc@47 339 Constructive logic lets us define all of the logical connectives in an aesthetically-appealing way, with orthogonal inductive definitions. That is, each connective is defined independently using a simple, shared mechanism. Constructivity also enables a trick called %\textit{%#<i>#program extraction#</i>#%}%, where we write programs by phrasing them as theorems to be proved. Since our proofs are just functional programs, we can extract executable programs from our final proofs, which we could not do as naturally with classical proofs.
adamc@47 340
adamc@47 341 We will see more about Coq's program extraction facility in a later chapter. However, I think it is worth interjecting another warning at this point, following up on the prior warning about taking the Curry-Howard correspondence too literally. It is possible to write programs by theorem-proving methods in Coq, but hardly anyone does it. It is almost always most useful to maintain the distinction between programs and proofs. If you write a program by proving a theorem, you are likely to run into algorithmic inefficiencies that you introduced in your proof to make it easier to prove. It is a shame to have to worry about such situations while proving tricky theorems, and it is a happy state of affairs that you almost certainly will not need to, with the ideal of extracting programs from proofs being confined mostly to theoretical studies. *)
adamc@48 342
adamc@48 343
adamc@48 344 (** * First-Order Logic *)
adamc@48 345
adamc@48 346 (** The [forall] connective of first-order logic, which we have seen in many examples so far, is built into Coq. Getting ahead of ourselves a bit, we can see it as the dependent function type constructor. In fact, implication and universal quantification are just different syntactic shorthands for the same Coq mechanism. A formula [P -> Q] is equivalent to [forall x : P, Q], where [x] does not appear in [Q]. That is, the "real" type of the implication says "for every proof of [P], there exists a proof of [Q]."
adamc@48 347
adamc@48 348 Existential quantification is defined in the standard library. *)
adamc@48 349
adamc@48 350 Print ex.
adamc@48 351 (** [[
adamc@48 352
adamc@48 353 Inductive ex (A : Type) (P : A -> Prop) : Prop :=
adamc@48 354 ex_intro : forall x : A, P x -> ex P
adamc@48 355 ]] *)
adamc@48 356
adamc@48 357 (** [ex] is parameterized by the type [A] that we quantify over, and by a predicate [P] over [A]s. We prove an existential by exhibiting some [x] of type [A], along with a proof of [P x]. As usual, there are tactics that save us from worrying about the low-level details most of the time. *)
adamc@48 358
adamc@48 359 Theorem exist1 : exists x : nat, x + 1 = 2.
adamc@55 360 (* begin thide *)
adamc@48 361 (** remove printing exists*)
adamc@55 362 (** We can start this proof with a tactic [exists], which should not be confused with the formula constructor shorthand of the same name. (In the PDF version of this document, the reverse 'E' appears instead of the text "exists" in formulas.) *)
adamc@48 363 exists 1.
adamc@48 364
adamc@48 365 (** The conclusion is replaced with a version using the existential witness that we announced. *)
adamc@48 366
adamc@48 367 (** [[
adamc@48 368
adamc@48 369 ============================
adamc@48 370 1 + 1 = 2
adamc@48 371 ]] *)
adamc@48 372
adamc@48 373 reflexivity.
adamc@55 374 (* end thide *)
adamc@48 375 Qed.
adamc@48 376
adamc@48 377 (** printing exists $\exists$ *)
adamc@48 378
adamc@48 379 (** We can also use tactics to reason about existential hypotheses. *)
adamc@48 380
adamc@48 381 Theorem exist2 : forall n m : nat, (exists x : nat, n + x = m) -> n <= m.
adamc@55 382 (* begin thide *)
adamc@48 383 (** We start by case analysis on the proof of the existential fact. *)
adamc@48 384 destruct 1.
adamc@48 385 (** [[
adamc@48 386
adamc@48 387 n : nat
adamc@48 388 m : nat
adamc@48 389 x : nat
adamc@48 390 H : n + x = m
adamc@48 391 ============================
adamc@48 392 n <= m
adamc@48 393 ]] *)
adamc@48 394
adamc@48 395 (** The goal has been replaced by a form where there is a new free variable [x], and where we have a new hypothesis that the body of the existential holds with [x] substituted for the old bound variable. From here, the proof is just about arithmetic and is easy to automate. *)
adamc@48 396
adamc@48 397 crush.
adamc@55 398 (* end thide *)
adamc@48 399 Qed.
adamc@48 400
adamc@48 401
adamc@48 402 (* begin hide *)
adamc@48 403 (* In-class exercises *)
adamc@48 404
adamc@48 405 Theorem forall_exists_commute : forall (A B : Type) (P : A -> B -> Prop),
adamc@48 406 (exists x : A, forall y : B, P x y) -> (forall y : B, exists x : A, P x y).
adamc@52 407 (* begin thide *)
adamc@52 408 intros.
adamc@52 409 destruct H.
adamc@52 410 exists x.
adamc@52 411 apply H.
adamc@52 412 (* end thide *)
adamc@48 413 Admitted.
adamc@48 414
adamc@48 415 (* end hide *)
adamc@48 416
adamc@48 417
adamc@48 418 (** The tactic [intuition] has a first-order cousin called [firstorder]. [firstorder] proves many formulas when only first-order reasoning is needed, and it tries to perform first-order simplifications in any case. First-order reasoning is much harder than propositional reasoning, so [firstorder] is much more likely than [intuition] to get stuck in a way that makes it run for long enough to be useless. *)
adamc@49 419
adamc@49 420
adamc@49 421 (** * Predicates with Implicit Equality *)
adamc@49 422
adamc@49 423 (** We start our exploration of a more complicated class of predicates with a simple example: an alternative way of characterizing when a natural number is zero. *)
adamc@49 424
adamc@49 425 Inductive isZero : nat -> Prop :=
adamc@49 426 | IsZero : isZero 0.
adamc@49 427
adamc@49 428 Theorem isZero_zero : isZero 0.
adamc@55 429 (* begin thide *)
adamc@49 430 constructor.
adamc@55 431 (* end thide *)
adamc@49 432 Qed.
adamc@49 433
adamc@49 434 (** We can call [isZero] a %\textit{%#<i>#judgment#</i>#%}%, in the sense often used in the semantics of programming languages. Judgments are typically defined in the style of %\textit{%#<i>#natural deduction#</i>#%}%, where we write a number of %\textit{%#<i>#inference rules#</i>#%}% with premises appearing above a solid line and a conclusion appearing below the line. In this example, the sole constructor [IsZero] of [isZero] can be thought of as the single inference rule for deducing [isZero], with nothing above the line and [isZero 0] below it. The proof of [isZero_zero] demonstrates how we can apply an inference rule.
adamc@49 435
adamc@49 436 The definition of [isZero] differs in an important way from all of the other inductive definitions that we have seen in this and the previous chapter. Instead of writing just [Set] or [Prop] after the colon, here we write [nat -> Prop]. We saw examples of parameterized types like [list], but there the parameters appeared with names %\textit{%#<i>#before#</i>#%}% the colon. Every constructor of a parameterized inductive type must have a range type that uses the same parameter, whereas the form we use here enables us to use different arguments to the type for different constructors.
adamc@49 437
adamc@49 438 For instance, [isZero] forces its argument to be [0]. We can see that the concept of equality is somehow implicit in the inductive definition mechanism. The way this is accomplished is similar to the way that logic variables are used in Prolog, and it is a very powerful mechanism that forms a foundation for formalizing all of mathematics. In fact, though it is natural to think of inductive types as folding in the functionality of equality, in Coq, the true situation is reversed, with equality defined as just another inductive type! *)
adamc@49 439
adamc@49 440 Print eq.
adamc@49 441 (** [[
adamc@49 442
adamc@49 443 Inductive eq (A : Type) (x : A) : A -> Prop := refl_equal : x = x
adamc@49 444 ]] *)
adamc@49 445
adamc@49 446 (** [eq] is the type we get behind the scenes when uses of infix [=] are expanded. We see that [eq] has both a parameter [x] that is fixed and an extra unnamed argument of the same type. It is crucial that the second argument is untied to the parameter in the type of [eq]; otherwise, we would have to prove that two values are equal even to be able to state the possibility of equality, which would more or less defeat the purpose of having an equality proposition. However, examining the type of equality's sole constructor [refl_equal], we see that we can only %\textit{%#<i>#prove#</i>#%}% equality when its two arguments are syntactically equal. This definition turns out to capture all of the basic properties of equality, and the equality-manipulating tactics that we have seen so far, like [reflexivity] and [rewrite], are implemented treating [eq] as just another inductive type with a well-chosen definition.
adamc@49 447
adamc@49 448 Returning to the example of [isZero], we can see how to make use of hypotheses that use this predicate. *)
adamc@49 449
adamc@49 450 Theorem isZero_plus : forall n m : nat, isZero m -> n + m = n.
adamc@55 451 (* begin thide *)
adamc@49 452 (** We want to proceed by cases on the proof of the assumption about [isZero]. *)
adamc@49 453 destruct 1.
adamc@49 454 (** [[
adamc@49 455
adamc@49 456 n : nat
adamc@49 457 ============================
adamc@49 458 n + 0 = n
adamc@49 459 ]] *)
adamc@49 460
adamc@49 461 (** Since [isZero] has only one constructor, we are presented with only one subgoal. The argument [m] to [isZero] is replaced with that type's argument from the single constructor [IsZero]. From this point, the proof is trivial. *)
adamc@49 462
adamc@49 463 crush.
adamc@55 464 (* end thide *)
adamc@49 465 Qed.
adamc@49 466
adamc@49 467 (** Another example seems at first like it should admit an analogous proof, but in fact provides a demonstration of one of the most basic gotchas of Coq proving. *)
adamc@49 468
adamc@49 469 Theorem isZero_contra : isZero 1 -> False.
adamc@55 470 (* begin thide *)
adamc@49 471 (** Let us try a proof by cases on the assumption, as in the last proof. *)
adamc@49 472 destruct 1.
adamc@49 473 (** [[
adamc@49 474
adamc@49 475 ============================
adamc@49 476 False
adamc@49 477 ]] *)
adamc@49 478
adamc@49 479 (** It seems that case analysis has not helped us much at all! Our sole hypothesis disappears, leaving us, if anything, worse off than we were before. What went wrong? We have met an important restriction in tactics like [destruct] and [induction] when applied to types with arguments. If the arguments are not already free variables, they will be replaced by new free variables internally before doing the case analysis or induction. Since the argument [1] to [isZero] is replaced by a fresh variable, we lose the crucial fact that it is not equal to [0].
adamc@49 480
adamc@49 481 Why does Coq use this restriction? We will discuss the issue in detail in a future chapter, when we see the dependently-typed programming techniques that would allow us to write this proof term manually. For now, we just say that the algorithmic problem of "logically complete case analysis" is undecidable when phrased in Coq's logic. A few tactics and design patterns that we will present in this chapter suffice in almost all cases. For the current example, what we want is a tactic called [inversion], which corresponds to the concept of inversion that is frequently used with natural deduction proof systems. *)
adamc@49 482
adamc@49 483 Undo.
adamc@49 484 inversion 1.
adamc@55 485 (* end thide *)
adamc@49 486 Qed.
adamc@49 487
adamc@49 488 (** What does [inversion] do? Think of it as a version of [destruct] that does its best to take advantage of the structure of arguments to inductive types. In this case, [inversion] completed the proof immediately, because it was able to detect that we were using [isZero] with an impossible argument.
adamc@49 489
adamc@49 490 Sometimes using [destruct] when you should have used [inversion] can lead to confusing results. To illustrate, consider an alternate proof attempt for the last theorem. *)
adamc@49 491
adamc@49 492 Theorem isZero_contra' : isZero 1 -> 2 + 2 = 5.
adamc@49 493 destruct 1.
adamc@49 494 (** [[
adamc@49 495
adamc@49 496 ============================
adamc@49 497 1 + 1 = 4
adamc@49 498 ]] *)
adamc@49 499
adamc@49 500 (** What on earth happened here? Internally, [destruct] replaced [1] with a fresh variable, and, trying to be helpful, it also replaced the occurrence of [1] within the unary representation of each number in the goal. This has the net effect of decrementing each of these numbers. If you are doing a proof and encounter a strange transmutation like this, there is a good chance that you should go back and replace a use of [destruct] with [inversion]. *)
adamc@49 501 Abort.
adamc@49 502
adamc@49 503
adamc@49 504 (* begin hide *)
adamc@49 505 (* In-class exercises *)
adamc@49 506
adamc@49 507 (* EX: Define an inductive type capturing when a list has exactly two elements. Prove that your predicate does not hold of the empty list, and prove that, whenever it holds of a list, the length of that list is two. *)
adamc@49 508
adamc@52 509 (* begin thide *)
adamc@52 510 Section twoEls.
adamc@52 511 Variable A : Type.
adamc@52 512
adamc@52 513 Inductive twoEls : list A -> Prop :=
adamc@52 514 | TwoEls : forall x y, twoEls (x :: y :: nil).
adamc@52 515
adamc@52 516 Theorem twoEls_nil : twoEls nil -> False.
adamc@52 517 inversion 1.
adamc@52 518 Qed.
adamc@52 519
adamc@52 520 Theorem twoEls_two : forall ls, twoEls ls -> length ls = 2.
adamc@52 521 inversion 1.
adamc@52 522 reflexivity.
adamc@52 523 Qed.
adamc@52 524 End twoEls.
adamc@52 525 (* end thide *)
adamc@52 526
adamc@49 527 (* end hide *)
adamc@49 528
adamc@50 529
adamc@50 530 (** * Recursive Predicates *)
adamc@50 531
adamc@50 532 (** We have already seen all of the ingredients we need to build interesting recursive predicates, like this predicate capturing even-ness. *)
adamc@50 533
adamc@50 534 Inductive even : nat -> Prop :=
adamc@50 535 | EvenO : even O
adamc@50 536 | EvenSS : forall n, even n -> even (S (S n)).
adamc@50 537
adamc@50 538 (** Think of [even] as another judgment defined by natural deduction rules. [EvenO] is a rule with nothing above the line and [even O] below the line, and [EvenSS] is a rule with [even n] above the line and [even (S (S n))] below.
adamc@50 539
adamc@50 540 The proof techniques of the last section are easily adapted. *)
adamc@50 541
adamc@50 542 Theorem even_0 : even 0.
adamc@55 543 (* begin thide *)
adamc@50 544 constructor.
adamc@55 545 (* end thide *)
adamc@50 546 Qed.
adamc@50 547
adamc@50 548 Theorem even_4 : even 4.
adamc@55 549 (* begin thide *)
adamc@50 550 constructor; constructor; constructor.
adamc@55 551 (* end thide *)
adamc@50 552 Qed.
adamc@50 553
adamc@50 554 (** It is not hard to see that sequences of constructor applications like the above can get tedious. We can avoid them using Coq's hint facility. *)
adamc@50 555
adamc@55 556 (* begin thide *)
adamc@50 557 Hint Constructors even.
adamc@50 558
adamc@50 559 Theorem even_4' : even 4.
adamc@50 560 auto.
adamc@50 561 Qed.
adamc@50 562
adamc@55 563 (* end thide *)
adamc@55 564
adamc@50 565 Theorem even_1_contra : even 1 -> False.
adamc@55 566 (* begin thide *)
adamc@50 567 inversion 1.
adamc@55 568 (* end thide *)
adamc@50 569 Qed.
adamc@50 570
adamc@50 571 Theorem even_3_contra : even 3 -> False.
adamc@55 572 (* begin thide *)
adamc@50 573 inversion 1.
adamc@50 574 (** [[
adamc@50 575
adamc@50 576 H : even 3
adamc@50 577 n : nat
adamc@50 578 H1 : even 1
adamc@50 579 H0 : n = 1
adamc@50 580 ============================
adamc@50 581 False
adamc@50 582 ]] *)
adamc@50 583
adamc@50 584 (** [inversion] can be a little overzealous at times, as we can see here with the introduction of the unused variable [n] and an equality hypothesis about it. For more complicated predicates, though, adding such assumptions is critical to dealing with the undecidability of general inversion. *)
adamc@50 585
adamc@50 586 inversion H1.
adamc@55 587 (* end thide *)
adamc@50 588 Qed.
adamc@50 589
adamc@50 590 (** We can also do inductive proofs about [even]. *)
adamc@50 591
adamc@50 592 Theorem even_plus : forall n m, even n -> even m -> even (n + m).
adamc@55 593 (* begin thide *)
adamc@50 594 (** It seems a reasonable first choice to proceed by induction on [n]. *)
adamc@50 595 induction n; crush.
adamc@50 596 (** [[
adamc@50 597
adamc@50 598 n : nat
adamc@50 599 IHn : forall m : nat, even n -> even m -> even (n + m)
adamc@50 600 m : nat
adamc@50 601 H : even (S n)
adamc@50 602 H0 : even m
adamc@50 603 ============================
adamc@50 604 even (S (n + m))
adamc@50 605 ]] *)
adamc@50 606
adamc@50 607 (** We will need to use the hypotheses [H] and [H0] somehow. The most natural choice is to invert [H]. *)
adamc@50 608
adamc@50 609 inversion H.
adamc@50 610 (** [[
adamc@50 611
adamc@50 612 n : nat
adamc@50 613 IHn : forall m : nat, even n -> even m -> even (n + m)
adamc@50 614 m : nat
adamc@50 615 H : even (S n)
adamc@50 616 H0 : even m
adamc@50 617 n0 : nat
adamc@50 618 H2 : even n0
adamc@50 619 H1 : S n0 = n
adamc@50 620 ============================
adamc@50 621 even (S (S n0 + m))
adamc@50 622 ]] *)
adamc@50 623
adamc@50 624 (** Simplifying the conclusion brings us to a point where we can apply a constructor. *)
adamc@50 625 simpl.
adamc@50 626 (** [[
adamc@50 627
adamc@50 628 ============================
adamc@50 629 even (S (S (n0 + m)))
adamc@50 630 ]] *)
adamc@50 631
adamc@50 632 constructor.
adamc@50 633 (** [[
adamc@50 634
adamc@50 635 ============================
adamc@50 636 even (n0 + m)
adamc@50 637 ]] *)
adamc@50 638
adamc@50 639 (** At this point, we would like to apply the inductive hypothesis, which is: *)
adamc@50 640 (** [[
adamc@50 641
adamc@50 642 IHn : forall m : nat, even n -> even m -> even (n + m)
adamc@50 643 ]] *)
adamc@50 644
adamc@50 645 (** Unfortunately, the goal mentions [n0] where it would need to mention [n] to match [IHn]. We could keep looking for a way to finish this proof from here, but it turns out that we can make our lives much easier by changing our basic strategy. Instead of inducting on the structure of [n], we should induct %\textit{%#<i>#on the structure of one of the [even] proofs#</i>#%}%. This technique is commonly called %\textit{%#<i>#rule induction#</i>#%}% in programming language semantics. In the setting of Coq, we have already seen how predicates are defined using the same inductive type mechanism as datatypes, so the fundamental unity of rule induction with "normal" induction is apparent. *)
adamc@50 646
adamc@50 647 Restart.
adamc@50 648
adamc@50 649 induction 1.
adamc@50 650 (** [[
adamc@50 651
adamc@50 652 m : nat
adamc@50 653 ============================
adamc@50 654 even m -> even (0 + m)
adamc@50 655
adamc@50 656 subgoal 2 is:
adamc@50 657 even m -> even (S (S n) + m)
adamc@50 658 ]] *)
adamc@50 659
adamc@50 660 (** The first case is easily discharged by [crush], based on the hint we added earlier to try the constructors of [even]. *)
adamc@50 661
adamc@50 662 crush.
adamc@50 663
adamc@50 664 (** Now we focus on the second case: *)
adamc@50 665 intro.
adamc@50 666
adamc@50 667 (** [[
adamc@50 668
adamc@50 669 m : nat
adamc@50 670 n : nat
adamc@50 671 H : even n
adamc@50 672 IHeven : even m -> even (n + m)
adamc@50 673 H0 : even m
adamc@50 674 ============================
adamc@50 675 even (S (S n) + m)
adamc@50 676 ]] *)
adamc@50 677
adamc@50 678 (** We simplify and apply a constructor, as in our last proof attempt. *)
adamc@50 679
adamc@50 680 simpl; constructor.
adamc@50 681 (** [[
adamc@50 682
adamc@50 683 ============================
adamc@50 684 even (n + m)
adamc@50 685 ]] *)
adamc@50 686
adamc@50 687 (** Now we have an exact match with our inductive hypothesis, and the remainder of the proof is trivial. *)
adamc@50 688
adamc@50 689 apply IHeven; assumption.
adamc@50 690
adamc@50 691 (** In fact, [crush] can handle all of the details of the proof once we declare the induction strategy. *)
adamc@50 692
adamc@50 693 Restart.
adamc@50 694 induction 1; crush.
adamc@55 695 (* end thide *)
adamc@50 696 Qed.
adamc@50 697
adamc@50 698 (** Induction on recursive predicates has similar pitfalls to those we encountered with inversion in the last section. *)
adamc@50 699
adamc@50 700 Theorem even_contra : forall n, even (S (n + n)) -> False.
adamc@55 701 (* begin thide *)
adamc@50 702 induction 1.
adamc@50 703 (** [[
adamc@50 704
adamc@50 705 n : nat
adamc@50 706 ============================
adamc@50 707 False
adamc@50 708
adamc@50 709 subgoal 2 is:
adamc@50 710 False
adamc@50 711 ]] *)
adamc@50 712
adamc@50 713 (** We are already sunk trying to prove the first subgoal, since the argument to [even] was replaced by a fresh variable internally. This time, we find it easiest to prove this theorem by way of a lemma. Instead of trusting [induction] to replace expressions with fresh variables, we do it ourselves, explicitly adding the appropriate equalities as new assumptions. *)
adamc@50 714 Abort.
adamc@50 715
adamc@50 716 Lemma even_contra' : forall n', even n' -> forall n, n' = S (n + n) -> False.
adamc@50 717 induction 1; crush.
adamc@50 718
adamc@54 719 (** At this point, it is useful to consider all cases of [n] and [n0] being zero or nonzero. Only one of these cases has any trickiness to it. *)
adamc@50 720 destruct n; destruct n0; crush.
adamc@50 721
adamc@50 722 (** [[
adamc@50 723
adamc@50 724 n : nat
adamc@50 725 H : even (S n)
adamc@50 726 IHeven : forall n0 : nat, S n = S (n0 + n0) -> False
adamc@50 727 n0 : nat
adamc@50 728 H0 : S n = n0 + S n0
adamc@50 729 ============================
adamc@50 730 False
adamc@50 731 ]] *)
adamc@50 732
adamc@50 733 (** At this point it is useful to use a theorem from the standard library, which we also proved with a different name in the last chapter. *)
adamc@50 734 Check plus_n_Sm.
adamc@50 735 (** [[
adamc@50 736
adamc@50 737 plus_n_Sm
adamc@50 738 : forall n m : nat, S (n + m) = n + S m
adamc@50 739 ]] *)
adamc@50 740
adamc@50 741 rewrite <- plus_n_Sm in H0.
adamc@50 742
adamc@50 743 (** The induction hypothesis lets us complete the proof. *)
adamc@50 744 apply IHeven with n0; assumption.
adamc@50 745
adamc@55 746 (** As usual, we can rewrite the proof to avoid referencing any locally-generated names, which makes our proof script more readable and more robust to changes in the theorem statement. *)
adamc@50 747 Restart.
adamc@50 748 Hint Rewrite <- plus_n_Sm : cpdt.
adamc@50 749
adamc@50 750 induction 1; crush;
adamc@50 751 match goal with
adamc@50 752 | [ H : S ?N = ?N0 + ?N0 |- _ ] => destruct N; destruct N0
adamc@50 753 end; crush; eauto.
adamc@50 754 Qed.
adamc@50 755
adamc@50 756 (** We write the proof in a way that avoids the use of local variable or hypothesis names, using the [match] tactic form to do pattern-matching on the goal. We use unification variables prefixed by question marks in the pattern, and we take advantage of the possibility to mention a unification variable twice in one pattern, to enforce equality between occurrences. The hint to rewrite with [plus_n_Sm] in a particular direction saves us from having to figure out the right place to apply that theorem, and we also take critical advantage of a new tactic, [eauto].
adamc@50 757
adamc@55 758 [crush] uses the tactic [intuition], which, when it runs out of tricks to try using only propositional logic, by default tries the tactic [auto], which we saw in an earlier example. [auto] attempts Prolog-style logic programming, searching through all proof trees up to a certain depth that are built only out of hints that have been registered with [Hint] commands. Compared to Prolog, [auto] places an important restriction: it never introduces new unification variables during search. That is, every time a rule is applied during proof search, all of its arguments must be deducible by studying the form of the goal. [eauto] relaxes this restriction, at the cost of possibly exponentially greater running time. In this particular case, we know that [eauto] has only a small space of proofs to search, so it makes sense to run it. It is common in effectively-automated Coq proofs to see a bag of standard tactics applied to pick off the "easy" subgoals, finishing with [eauto] to handle the tricky parts that can benefit from ad-hoc exhaustive search.
adamc@50 759
adamc@50 760 The original theorem now follows trivially from our lemma. *)
adamc@50 761
adamc@50 762 Theorem even_contra : forall n, even (S (n + n)) -> False.
adamc@52 763 intros; eapply even_contra'; eauto.
adamc@50 764 Qed.
adamc@52 765
adamc@52 766 (** We use a variant [eapply] of [apply] which has the same relationship to [apply] as [eauto] has to [auto]. [apply] only succeeds if all arguments to the rule being used can be determined from the form of the goal, whereas [eapply] will introduce unification variables for undetermined arguments. [eauto] is able to determine the right values for those unification variables.
adamc@52 767
adamc@52 768 By considering an alternate attempt at proving the lemma, we can see another common pitfall of inductive proofs in Coq. Imagine that we had tried to prove [even_contra'] with all of the [forall] quantifiers moved to the front of the lemma statement. *)
adamc@52 769
adamc@52 770 Lemma even_contra'' : forall n' n, even n' -> n' = S (n + n) -> False.
adamc@52 771 induction 1; crush;
adamc@52 772 match goal with
adamc@52 773 | [ H : S ?N = ?N0 + ?N0 |- _ ] => destruct N; destruct N0
adamc@52 774 end; crush; eauto.
adamc@52 775
adamc@52 776 (** One subgoal remains: *)
adamc@52 777
adamc@52 778 (** [[
adamc@52 779
adamc@52 780 n : nat
adamc@52 781 H : even (S (n + n))
adamc@52 782 IHeven : S (n + n) = S (S (S (n + n))) -> False
adamc@52 783 ============================
adamc@52 784 False
adamc@52 785 ]] *)
adamc@52 786
adamc@52 787 (** We are out of luck here. The inductive hypothesis is trivially true, since its assumption is false. In the version of this proof that succeeded, [IHeven] had an explicit quantification over [n]. This is because the quantification of [n] %\textit{%#<i>#appeared after the thing we are inducting on#</i>#%}% in the theorem statement. In general, quantified variables and hypotheses that appear before the induction object in the theorem statement stay fixed throughout the inductive proof. Variables and hypotheses that are quantified after the induction object may be varied explicitly in uses of inductive hypotheses.
adamc@52 788
adamc@52 789 Why should Coq implement [induction] this way? One answer is that it avoids burdening this basic tactic with additional heuristic smarts, but that is not the whole picture. Imagine that [induction] analyzed dependencies among variables and reordered quantifiers to preserve as much freedom as possible in later uses of inductive hypotheses. This could make the inductive hypotheses more complex, which could in turn cause particular automation machinery to fail when it would have succeeded before. In general, we want to avoid quantifiers in our proofs whenever we can, and that goal is furthered by the refactoring that the [induction] tactic forces us to do. *)
adamc@55 790 (* end thide *)
adamc@51 791 Abort.
adamc@51 792
adamc@52 793
adamc@52 794 (* begin hide *)
adamc@52 795 (* In-class exercises *)
adamc@52 796
adamc@52 797 (* EX: Define a type [prop] of simple boolean formulas made up only of truth, falsehood, binary conjunction, and binary disjunction. Define an inductive predicate [holds] that captures when [prop]s are valid, and define a predicate [falseFree] that captures when a [prop] does not contain the "false" formula. Prove that every false-free [prop] is valid. *)
adamc@52 798
adamc@52 799 (* begin thide *)
adamc@52 800 Inductive prop : Set :=
adamc@52 801 | Tru : prop
adamc@52 802 | Fals : prop
adamc@52 803 | And : prop -> prop -> prop
adamc@52 804 | Or : prop -> prop -> prop.
adamc@52 805
adamc@52 806 Inductive holds : prop -> Prop :=
adamc@52 807 | HTru : holds Tru
adamc@52 808 | HAnd : forall p1 p2, holds p1 -> holds p2 -> holds (And p1 p2)
adamc@52 809 | HOr1 : forall p1 p2, holds p1 -> holds (Or p1 p2)
adamc@52 810 | HOr2 : forall p1 p2, holds p2 -> holds (Or p1 p2).
adamc@52 811
adamc@52 812 Inductive falseFree : prop -> Prop :=
adamc@52 813 | FFTru : falseFree Tru
adamc@52 814 | FFAnd : forall p1 p2, falseFree p1 -> falseFree p2 -> falseFree (And p1 p2)
adamc@52 815 | FFNot : forall p1 p2, falseFree p1 -> falseFree p2 -> falseFree (Or p1 p2).
adamc@52 816
adamc@52 817 Hint Constructors holds.
adamc@52 818
adamc@52 819 Theorem falseFree_holds : forall p, falseFree p -> holds p.
adamc@52 820 induction 1; crush.
adamc@52 821 Qed.
adamc@52 822 (* end thide *)
adamc@52 823
adamc@52 824
adamc@52 825 (* EX: Define an inductive type [prop'] that is the same as [prop] but omits the possibility for falsehood. Define a proposition [holds'] for [prop'] that is analogous to [holds]. Define a function [propify] for translating [prop']s to [prop]s. Prove that, for any [prop'] [p], if [propify p] is valid, then so is [p]. *)
adamc@52 826
adamc@52 827 (* begin thide *)
adamc@52 828 Inductive prop' : Set :=
adamc@52 829 | Tru' : prop'
adamc@52 830 | And' : prop' -> prop' -> prop'
adamc@52 831 | Or' : prop' -> prop' -> prop'.
adamc@52 832
adamc@52 833 Inductive holds' : prop' -> Prop :=
adamc@52 834 | HTru' : holds' Tru'
adamc@52 835 | HAnd' : forall p1 p2, holds' p1 -> holds' p2 -> holds' (And' p1 p2)
adamc@52 836 | HOr1' : forall p1 p2, holds' p1 -> holds' (Or' p1 p2)
adamc@52 837 | HOr2' : forall p1 p2, holds' p2 -> holds' (Or' p1 p2).
adamc@52 838
adamc@52 839 Fixpoint propify (p : prop') : prop :=
adamc@52 840 match p with
adamc@52 841 | Tru' => Tru
adamc@52 842 | And' p1 p2 => And (propify p1) (propify p2)
adamc@52 843 | Or' p1 p2 => Or (propify p1) (propify p2)
adamc@52 844 end.
adamc@52 845
adamc@52 846 Hint Constructors holds'.
adamc@52 847
adamc@52 848 Lemma propify_holds' : forall p', holds p' -> forall p, p' = propify p -> holds' p.
adamc@52 849 induction 1; crush; destruct p; crush.
adamc@52 850 Qed.
adamc@52 851
adamc@52 852 Theorem propify_holds : forall p, holds (propify p) -> holds' p.
adamc@52 853 intros; eapply propify_holds'; eauto.
adamc@52 854 Qed.
adamc@52 855 (* end thide *)
adamc@52 856
adamc@52 857 (* end hide *)