Mercurial > cpdt > repo

annotate src/Universes.v @ 564:5504235ea06d

Find changesets by keywords (author, files, the commit message), revision number or hash, or revset expression.
Link to RIT CSCI 740
author Adam Chlipala <adam@chlipala.net>
date Sun, 25 Mar 2018 13:03:34 -0400
parents af97676583f3
children 81d63d9c1cc5
rev   line source
adam@534 1 (* Copyright (c) 2009-2012, 2015, Adam Chlipala
adamc@227 2 *
adamc@227 3 * This work is licensed under a
adamc@227 4 * Creative Commons Attribution-Noncommercial-No Derivative Works 3.0
adamc@227 5 * Unported License.
adamc@227 6 * The license text is available at:
adamc@227 7 * http://creativecommons.org/licenses/by-nc-nd/3.0/
adamc@227 8 *)
adamc@227 9
adamc@227 10 (* begin hide *)
adam@377 11 Require Import List.
adam@377 12
adam@534 13 Require Import DepList Cpdt.CpdtTactics.
adamc@227 14
adam@563 15 Require Extraction.
adam@563 16
adamc@227 17 Set Implicit Arguments.
adam@534 18 Set Asymmetric Patterns.
adamc@227 19 (* end hide *)
adamc@227 20
adam@398 21 (** printing $ %({}*% #(<a/>*# *)
adam@398 22 (** printing ^ %*{})% #*<a/>)# *)
adam@398 23
adam@398 24
adamc@227 25
adamc@227 26 (** %\chapter{Universes and Axioms}% *)
adamc@227 27
adam@343 28 (** Many traditional theorems can be proved in Coq without special knowledge of CIC, the logic behind the prover. A development just seems to be using a particular ASCII notation for standard formulas based on %\index{set theory}%set theory. Nonetheless, as we saw in Chapter 4, CIC differs from set theory in starting from fewer orthogonal primitives. It is possible to define the usual logical connectives as derived notions. The foundation of it all is a dependently typed functional programming language, based on dependent function types and inductive type families. By using the facilities of this language directly, we can accomplish some things much more easily than in mainstream math.
adamc@227 29
adam@343 30 %\index{Gallina}%Gallina, which adds features to the more theoretical CIC%~\cite{CIC}%, is the logic implemented in Coq. It has a relatively simple foundation that can be defined rigorously in a page or two of formal proof rules. Still, there are some important subtleties that have practical ramifications. This chapter focuses on those subtleties, avoiding formal metatheory in favor of example code. *)
adamc@227 31
adamc@227 32
adamc@227 33 (** * The [Type] Hierarchy *)
adamc@227 34
adam@343 35 (** %\index{type hierarchy}%Every object in Gallina has a type. *)
adamc@227 36
adamc@227 37 Check 0.
adamc@227 38 (** %\vspace{-.15in}% [[
adamc@227 39 0
adamc@227 40 : nat
adamc@227 41 ]]
adamc@227 42
adamc@227 43 It is natural enough that zero be considered as a natural number. *)
adamc@227 44
adamc@227 45 Check nat.
adamc@227 46 (** %\vspace{-.15in}% [[
adamc@227 47 nat
adamc@227 48 : Set
adamc@227 49 ]]
adamc@227 50
adam@429 51 From a set theory perspective, it is unsurprising to consider the natural numbers as a "set." *)
adamc@227 52
adamc@227 53 Check Set.
adamc@227 54 (** %\vspace{-.15in}% [[
adamc@227 55 Set
adamc@227 56 : Type
adamc@227 57 ]]
adamc@227 58
adam@409 59 The type [Set] may be considered as the set of all sets, a concept that set theory handles in terms of%\index{class (in set theory)}% _classes_. In Coq, this more general notion is [Type]. *)
adamc@227 60
adamc@227 61 Check Type.
adamc@227 62 (** %\vspace{-.15in}% [[
adamc@227 63 Type
adamc@227 64 : Type
adamc@227 65 ]]
adamc@227 66
adam@429 67 Strangely enough, [Type] appears to be its own type. It is known that polymorphic languages with this property are inconsistent, via %\index{Girard's paradox}%Girard's paradox%~\cite{GirardsParadox}%. That is, using such a language to encode proofs is unwise, because it is possible to "prove" any proposition. What is really going on here?
adamc@227 68
adam@343 69 Let us repeat some of our queries after toggling a flag related to Coq's printing behavior.%\index{Vernacular commands!Set Printing Universes}% *)
adamc@227 70
adamc@227 71 Set Printing Universes.
adamc@227 72
adamc@227 73 Check nat.
adamc@227 74 (** %\vspace{-.15in}% [[
adamc@227 75 nat
adamc@227 76 : Set
adam@302 77 ]]
adam@398 78 *)
adamc@227 79
adamc@227 80 Check Set.
adamc@227 81 (** %\vspace{-.15in}% [[
adamc@227 82 Set
adamc@227 83 : Type $ (0)+1 ^
adam@302 84 ]]
adam@302 85 *)
adamc@227 86
adamc@227 87 Check Type.
adamc@227 88 (** %\vspace{-.15in}% [[
adamc@227 89 Type $ Top.3 ^
adamc@227 90 : Type $ (Top.3)+1 ^
adamc@227 91 ]]
adamc@227 92
adam@429 93 Occurrences of [Type] are annotated with some additional information, inside comments. These annotations have to do with the secret behind [Type]: it really stands for an infinite hierarchy of types. The type of [Set] is [Type(0)], the type of [Type(0)] is [Type(1)], the type of [Type(1)] is [Type(2)], and so on. This is how we avoid the "[Type : Type]" paradox. As a convenience, the universe hierarchy drives Coq's one variety of subtyping. Any term whose type is [Type] at level [i] is automatically also described by [Type] at level [j] when [j > i].
adamc@227 94
adam@398 95 In the outputs of our first [Check] query, we see that the type level of [Set]'s type is [(0)+1]. Here [0] stands for the level of [Set], and we increment it to arrive at the level that _classifies_ [Set].
adamc@227 96
adam@488 97 In the third query's output, we see that the occurrence of [Type] that we check is assigned a fresh%\index{universe variable}% _universe variable_ [Top.3]. The output type increments [Top.3] to move up a level in the universe hierarchy. As we write code that uses definitions whose types mention universe variables, unification may refine the values of those variables. Luckily, the user rarely has to worry about the details.
adamc@227 98
adam@409 99 Another crucial concept in CIC is%\index{predicativity}% _predicativity_. Consider these queries. *)
adamc@227 100
adamc@227 101 Check forall T : nat, fin T.
adamc@227 102 (** %\vspace{-.15in}% [[
adamc@227 103 forall T : nat, fin T
adamc@227 104 : Set
adam@302 105 ]]
adam@302 106 *)
adamc@227 107
adamc@227 108 Check forall T : Set, T.
adamc@227 109 (** %\vspace{-.15in}% [[
adamc@227 110 forall T : Set, T
adamc@227 111 : Type $ max(0, (0)+1) ^
adam@302 112 ]]
adam@302 113 *)
adamc@227 114
adamc@227 115 Check forall T : Type, T.
adamc@227 116 (** %\vspace{-.15in}% [[
adamc@227 117 forall T : Type $ Top.9 ^ , T
adamc@227 118 : Type $ max(Top.9, (Top.9)+1) ^
adamc@227 119 ]]
adamc@227 120
adamc@227 121 These outputs demonstrate the rule for determining which universe a [forall] type lives in. In particular, for a type [forall x : T1, T2], we take the maximum of the universes of [T1] and [T2]. In the first example query, both [T1] ([nat]) and [T2] ([fin T]) are in [Set], so the [forall] type is in [Set], too. In the second query, [T1] is [Set], which is at level [(0)+1]; and [T2] is [T], which is at level [0]. Thus, the [forall] exists at the maximum of these two levels. The third example illustrates the same outcome, where we replace [Set] with an occurrence of [Type] that is assigned universe variable [Top.9]. This universe variable appears in the places where [0] appeared in the previous query.
adamc@227 122
adam@287 123 The behind-the-scenes manipulation of universe variables gives us predicativity. Consider this simple definition of a polymorphic identity function, where the first argument [T] will automatically be marked as implicit, since it can be inferred from the type of the second argument [x]. *)
adamc@227 124
adamc@227 125 Definition id (T : Set) (x : T) : T := x.
adamc@227 126
adamc@227 127 Check id 0.
adamc@227 128 (** %\vspace{-.15in}% [[
adamc@227 129 id 0
adamc@227 130 : nat
adamc@227 131
adamc@227 132 Check id Set.
adam@343 133 ]]
adamc@227 134
adam@343 135 <<
adamc@227 136 Error: Illegal application (Type Error):
adamc@227 137 ...
adam@479 138 The 1st term has type "Type (* (Top.15)+1 *)"
adam@479 139 which should be coercible to "Set".
adam@343 140 >>
adamc@227 141
adam@343 142 The parameter [T] of [id] must be instantiated with a [Set]. The type [nat] is a [Set], but [Set] is not. We can try fixing the problem by generalizing our definition of [id]. *)
adamc@227 143
adamc@227 144 Reset id.
adamc@227 145 Definition id (T : Type) (x : T) : T := x.
adamc@227 146 Check id 0.
adamc@227 147 (** %\vspace{-.15in}% [[
adamc@227 148 id 0
adamc@227 149 : nat
adam@302 150 ]]
adam@302 151 *)
adamc@227 152
adamc@227 153 Check id Set.
adamc@227 154 (** %\vspace{-.15in}% [[
adamc@227 155 id Set
adamc@227 156 : Type $ Top.17 ^
adam@302 157 ]]
adam@302 158 *)
adamc@227 159
adamc@227 160 Check id Type.
adamc@227 161 (** %\vspace{-.15in}% [[
adamc@227 162 id Type $ Top.18 ^
adamc@227 163 : Type $ Top.19 ^
adam@302 164 ]]
adam@302 165 *)
adamc@227 166
adamc@227 167 (** So far so good. As we apply [id] to different [T] values, the inferred index for [T]'s [Type] occurrence automatically moves higher up the type hierarchy.
adamc@227 168 [[
adamc@227 169 Check id id.
adam@343 170 ]]
adamc@227 171
adam@343 172 <<
adamc@227 173 Error: Universe inconsistency (cannot enforce Top.16 < Top.16).
adam@343 174 >>
adamc@227 175
adam@479 176 %\index{universe inconsistency}%This error message reminds us that the universe variable for [T] still exists, even though it is usually hidden. To apply [id] to itself, that variable would need to be less than itself in the type hierarchy. Universe inconsistency error messages announce cases like this one where a term could only type-check by violating an implied constraint over universe variables. Such errors demonstrate that [Type] is _predicative_, where this word has a CIC meaning closely related to its usual mathematical meaning. A predicative system enforces the constraint that, when an object is defined using some sort of quantifier, none of the quantifiers may ever be instantiated with the object itself. %\index{impredicativity}%Impredicativity is associated with popular paradoxes in set theory, involving inconsistent constructions like "the set of all sets that do not contain themselves" (%\index{Russell's paradox}%Russell's paradox). Similar paradoxes would result from uncontrolled impredicativity in Coq. *)
adamc@227 177
adamc@227 178
adamc@227 179 (** ** Inductive Definitions *)
adamc@227 180
adam@505 181 (** Predicativity restrictions also apply to inductive definitions. As an example, let us consider a type of expression trees that allows injection of any native Coq value. The idea is that an [exp T] stands for an encoded expression of type [T].
adamc@227 182 [[
adamc@227 183 Inductive exp : Set -> Set :=
adamc@227 184 | Const : forall T : Set, T -> exp T
adamc@227 185 | Pair : forall T1 T2, exp T1 -> exp T2 -> exp (T1 * T2)
adamc@227 186 | Eq : forall T, exp T -> exp T -> exp bool.
adam@343 187 ]]
adamc@227 188
adam@343 189 <<
adamc@227 190 Error: Large non-propositional inductive types must be in Type.
adam@343 191 >>
adamc@227 192
adam@409 193 This definition is%\index{large inductive types}% _large_ in the sense that at least one of its constructors takes an argument whose type has type [Type]. Coq would be inconsistent if we allowed definitions like this one in their full generality. Instead, we must change [exp] to live in [Type]. We will go even further and move [exp]'s index to [Type] as well. *)
adamc@227 194
adamc@227 195 Inductive exp : Type -> Type :=
adamc@227 196 | Const : forall T, T -> exp T
adamc@227 197 | Pair : forall T1 T2, exp T1 -> exp T2 -> exp (T1 * T2)
adamc@227 198 | Eq : forall T, exp T -> exp T -> exp bool.
adamc@227 199
adam@505 200 (** Note that before we had to include an annotation [: Set] for the variable [T] in [Const]'s type, but we need no annotation now. When the type of a variable is not known, and when that variable is used in a context where only types are allowed, Coq infers that the variable is of type [Type], the right behavior here, though it was wrong for the [Set] version of [exp].
adamc@228 201
adamc@228 202 Our new definition is accepted. We can build some sample expressions. *)
adamc@227 203
adamc@227 204 Check Const 0.
adamc@227 205 (** %\vspace{-.15in}% [[
adamc@227 206 Const 0
adamc@227 207 : exp nat
adam@302 208 ]]
adam@302 209 *)
adamc@227 210
adamc@227 211 Check Pair (Const 0) (Const tt).
adamc@227 212 (** %\vspace{-.15in}% [[
adamc@227 213 Pair (Const 0) (Const tt)
adamc@227 214 : exp (nat * unit)
adam@302 215 ]]
adam@302 216 *)
adamc@227 217
adamc@227 218 Check Eq (Const Set) (Const Type).
adamc@227 219 (** %\vspace{-.15in}% [[
adamc@228 220 Eq (Const Set) (Const Type $ Top.59 ^ )
adamc@227 221 : exp bool
adamc@227 222 ]]
adamc@227 223
adamc@227 224 We can check many expressions, including fancy expressions that include types. However, it is not hard to hit a type-checking wall.
adamc@227 225 [[
adamc@227 226 Check Const (Const O).
adam@343 227 ]]
adamc@227 228
adam@343 229 <<
adamc@227 230 Error: Universe inconsistency (cannot enforce Top.42 < Top.42).
adam@343 231 >>
adamc@227 232
adamc@227 233 We are unable to instantiate the parameter [T] of [Const] with an [exp] type. To see why, it is helpful to print the annotated version of [exp]'s inductive definition. *)
adam@417 234 (** [[
adamc@227 235 Print exp.
adam@417 236 ]]
adam@444 237 %\vspace{-.15in}%[[
adamc@227 238 Inductive exp
adamc@227 239 : Type $ Top.8 ^ ->
adamc@227 240 Type
adamc@227 241 $ max(0, (Top.11)+1, (Top.14)+1, (Top.15)+1, (Top.19)+1) ^ :=
adamc@227 242 Const : forall T : Type $ Top.11 ^ , T -> exp T
adamc@227 243 | Pair : forall (T1 : Type $ Top.14 ^ ) (T2 : Type $ Top.15 ^ ),
adamc@227 244 exp T1 -> exp T2 -> exp (T1 * T2)
adamc@227 245 | Eq : forall T : Type $ Top.19 ^ , exp T -> exp T -> exp bool
adamc@227 246 ]]
adamc@227 247
adam@505 248 We see that the index type of [exp] has been assigned to universe level [Top.8]. In addition, each of the four occurrences of [Type] in the types of the constructors gets its own universe variable. Each of these variables appears explicitly in the type of [exp]. In particular, any type [exp T] lives at a universe level found by incrementing by one the maximum of the four argument variables. Therefore, [exp] _must_ live at a higher universe level than any type which may be passed to one of its constructors. This consequence led to the universe inconsistency.
adamc@227 249
adam@429 250 Strangely, the universe variable [Top.8] only appears in one place. Is there no restriction imposed on which types are valid arguments to [exp]? In fact, there is a restriction, but it only appears in a global set of universe constraints that are maintained "off to the side," not appearing explicitly in types. We can print the current database.%\index{Vernacular commands!Print Universes}% *)
adamc@227 251
adamc@227 252 Print Universes.
adamc@227 253 (** %\vspace{-.15in}% [[
adamc@227 254 Top.19 < Top.9 <= Top.8
adamc@227 255 Top.15 < Top.9 <= Top.8 <= Coq.Init.Datatypes.38
adamc@227 256 Top.14 < Top.9 <= Top.8 <= Coq.Init.Datatypes.37
adamc@227 257 Top.11 < Top.9 <= Top.8
adamc@227 258 ]]
adamc@227 259
adam@343 260 The command outputs many more constraints, but we have collected only those that mention [Top] variables. We see one constraint for each universe variable associated with a constructor argument from [exp]'s definition. Universe variable [Top.19] is the type argument to [Eq]. The constraint for [Top.19] effectively says that [Top.19] must be less than [Top.8], the universe of [exp]'s indices; an intermediate variable [Top.9] appears as an artifact of the way the constraint was generated.
adamc@227 261
adamc@227 262 The next constraint, for [Top.15], is more complicated. This is the universe of the second argument to the [Pair] constructor. Not only must [Top.15] be less than [Top.8], but it also comes out that [Top.8] must be less than [Coq.Init.Datatypes.38]. What is this new universe variable? It is from the definition of the [prod] inductive family, to which types of the form [A * B] are desugared. *)
adamc@227 263
adam@417 264 (* begin hide *)
adam@437 265 (* begin thide *)
adam@417 266 Inductive prod := pair.
adam@417 267 Reset prod.
adam@437 268 (* end thide *)
adam@417 269 (* end hide *)
adam@417 270
adam@444 271 (** %\vspace{-.3in}%[[
adamc@227 272 Print prod.
adam@417 273 ]]
adam@444 274 %\vspace{-.15in}%[[
adamc@227 275 Inductive prod (A : Type $ Coq.Init.Datatypes.37 ^ )
adamc@227 276 (B : Type $ Coq.Init.Datatypes.38 ^ )
adamc@227 277 : Type $ max(Coq.Init.Datatypes.37, Coq.Init.Datatypes.38) ^ :=
adamc@227 278 pair : A -> B -> A * B
adamc@227 279 ]]
adamc@227 280
adamc@227 281 We see that the constraint is enforcing that indices to [exp] must not live in a higher universe level than [B]-indices to [prod]. The next constraint above establishes a symmetric condition for [A].
adamc@227 282
adamc@227 283 Thus it is apparent that Coq maintains a tortuous set of universe variable inequalities behind the scenes. It may look like some functions are polymorphic in the universe levels of their arguments, but what is really happening is imperative updating of a system of constraints, such that all uses of a function are consistent with a global set of universe levels. When the constraint system may not be evolved soundly, we get a universe inconsistency error.
adamc@227 284
adamc@227 285 %\medskip%
adamc@227 286
adam@505 287 The annotated definition of [prod] reveals something interesting. A type [prod A B] lives at a universe that is the maximum of the universes of [A] and [B]. From our earlier experiments, we might expect that [prod]'s universe would in fact need to be _one higher_ than the maximum. The critical difference is that, in the definition of [prod], [A] and [B] are defined as _parameters_; that is, they appear named to the left of the main colon, rather than appearing (possibly unnamed) to the right.
adamc@227 288
adamc@231 289 Parameters are not as flexible as normal inductive type arguments. The range types of all of the constructors of a parameterized type must share the same parameters. Nonetheless, when it is possible to define a polymorphic type in this way, we gain the ability to use the new type family in more ways, without triggering universe inconsistencies. For instance, nested pairs of types are perfectly legal. *)
adamc@227 290
adamc@227 291 Check (nat, (Type, Set)).
adamc@227 292 (** %\vspace{-.15in}% [[
adamc@227 293 (nat, (Type $ Top.44 ^ , Set))
adamc@227 294 : Set * (Type $ Top.45 ^ * Type $ Top.46 ^ )
adamc@227 295 ]]
adamc@227 296
adamc@227 297 The same cannot be done with a counterpart to [prod] that does not use parameters. *)
adamc@227 298
adamc@227 299 Inductive prod' : Type -> Type -> Type :=
adamc@227 300 | pair' : forall A B : Type, A -> B -> prod' A B.
adam@444 301 (** %\vspace{-.15in}%[[
adamc@227 302 Check (pair' nat (pair' Type Set)).
adam@343 303 ]]
adamc@227 304
adam@343 305 <<
adamc@227 306 Error: Universe inconsistency (cannot enforce Top.51 < Top.51).
adam@343 307 >>
adamc@227 308
adamc@233 309 The key benefit parameters bring us is the ability to avoid quantifying over types in the types of constructors. Such quantification induces less-than constraints, while parameters only introduce less-than-or-equal-to constraints.
adamc@233 310
adam@343 311 Coq includes one more (potentially confusing) feature related to parameters. While Gallina does not support real %\index{universe polymorphism}%universe polymorphism, there is a convenience facility that mimics universe polymorphism in some cases. We can illustrate what this means with a simple example. *)
adamc@233 312
adamc@233 313 Inductive foo (A : Type) : Type :=
adamc@233 314 | Foo : A -> foo A.
adamc@229 315
adamc@229 316 (* begin hide *)
adamc@229 317 Unset Printing Universes.
adamc@229 318 (* end hide *)
adamc@229 319
adamc@233 320 Check foo nat.
adamc@233 321 (** %\vspace{-.15in}% [[
adamc@233 322 foo nat
adamc@233 323 : Set
adam@302 324 ]]
adam@302 325 *)
adamc@233 326
adamc@233 327 Check foo Set.
adamc@233 328 (** %\vspace{-.15in}% [[
adamc@233 329 foo Set
adamc@233 330 : Type
adam@302 331 ]]
adam@302 332 *)
adamc@233 333
adamc@233 334 Check foo True.
adamc@233 335 (** %\vspace{-.15in}% [[
adamc@233 336 foo True
adamc@233 337 : Prop
adamc@233 338 ]]
adamc@233 339
adam@429 340 The basic pattern here is that Coq is willing to automatically build a "copied-and-pasted" version of an inductive definition, where some occurrences of [Type] have been replaced by [Set] or [Prop]. In each context, the type-checker tries to find the valid replacements that are lowest in the type hierarchy. Automatic cloning of definitions can be much more convenient than manual cloning. We have already taken advantage of the fact that we may re-use the same families of tuple and list types to form values in [Set] and [Type].
adamc@233 341
adamc@233 342 Imitation polymorphism can be confusing in some contexts. For instance, it is what is responsible for this weird behavior. *)
adamc@233 343
adamc@233 344 Inductive bar : Type := Bar : bar.
adamc@233 345
adamc@233 346 Check bar.
adamc@233 347 (** %\vspace{-.15in}% [[
adamc@233 348 bar
adamc@233 349 : Prop
adamc@233 350 ]]
adamc@233 351
adamc@233 352 The type that Coq comes up with may be used in strictly more contexts than the type one might have expected. *)
adamc@233 353
adamc@229 354
adam@388 355 (** ** Deciphering Baffling Messages About Inability to Unify *)
adam@388 356
adam@388 357 (** One of the most confusing sorts of Coq error messages arises from an interplay between universes, syntax notations, and %\index{implicit arguments}%implicit arguments. Consider the following innocuous lemma, which is symmetry of equality for the special case of types. *)
adam@388 358
adam@388 359 Theorem symmetry : forall A B : Type,
adam@388 360 A = B
adam@388 361 -> B = A.
adam@388 362 intros ? ? H; rewrite H; reflexivity.
adam@388 363 Qed.
adam@388 364
adam@388 365 (** Let us attempt an admittedly silly proof of the following theorem. *)
adam@388 366
adam@388 367 Theorem illustrative_but_silly_detour : unit = unit.
adam@444 368 (** %\vspace{-.25in}%[[
adam@444 369 apply symmetry.
adam@388 370 ]]
adam@388 371 <<
adam@388 372 Error: Impossible to unify "?35 = ?34" with "unit = unit".
adam@388 373 >>
adam@388 374
adam@458 375 Coq tells us that we cannot, in fact, apply our lemma [symmetry] here, but the error message seems defective. In particular, one might think that [apply] should unify [?35] and [?34] with [unit] to ensure that the unification goes through. In fact, the issue is in a part of the unification problem that is _not_ shown to us in this error message!
adam@388 376
adam@510 377 The following command is the secret to getting better error messages in such cases:%\index{Vernacular commands!Set Printing All}% *)
adam@388 378
adam@388 379 Set Printing All.
adam@444 380 (** %\vspace{-.15in}%[[
adam@444 381 apply symmetry.
adam@388 382 ]]
adam@388 383 <<
adam@388 384 Error: Impossible to unify "@eq Type ?46 ?45" with "@eq Set unit unit".
adam@388 385 >>
adam@388 386
adam@398 387 Now we can see the problem: it is the first, _implicit_ argument to the underlying equality function [eq] that disagrees across the two terms. The universe [Set] may be both an element and a subtype of [Type], but the two are not definitionally equal. *)
adam@388 388
adam@388 389 Abort.
adam@388 390
adam@388 391 (** A variety of changes to the theorem statement would lead to use of [Type] as the implicit argument of [eq]. Here is one such change. *)
adam@388 392
adam@388 393 Theorem illustrative_but_silly_detour : (unit : Type) = unit.
adam@388 394 apply symmetry; reflexivity.
adam@388 395 Qed.
adam@388 396
adam@388 397 (** There are many related issues that can come up with error messages, where one or both of notations and implicit arguments hide important details. The [Set Printing All] command turns off all such features and exposes underlying CIC terms.
adam@388 398
adam@388 399 For completeness, we mention one other class of confusing error message about inability to unify two terms that look obviously unifiable. Each unification variable has a scope; a unification variable instantiation may not mention variables that were not already defined within that scope, at the point in proof search where the unification variable was introduced. Consider this illustrative example: *)
adam@388 400
adam@388 401 Unset Printing All.
adam@388 402
adam@388 403 Theorem ex_symmetry : (exists x, x = 0) -> (exists x, 0 = x).
adam@435 404 eexists.
adam@388 405 (** %\vspace{-.15in}%[[
adam@388 406 H : exists x : nat, x = 0
adam@388 407 ============================
adam@388 408 0 = ?98
adam@388 409 ]]
adam@388 410 *)
adam@388 411
adam@388 412 destruct H.
adam@388 413 (** %\vspace{-.15in}%[[
adam@388 414 x : nat
adam@388 415 H : x = 0
adam@388 416 ============================
adam@388 417 0 = ?99
adam@388 418 ]]
adam@388 419 *)
adam@388 420
adam@444 421 (** %\vspace{-.2in}%[[
adam@444 422 symmetry; exact H.
adam@388 423 ]]
adam@388 424
adam@388 425 <<
adam@388 426 Error: In environment
adam@388 427 x : nat
adam@388 428 H : x = 0
adam@388 429 The term "H" has type "x = 0" while it is expected to have type
adam@388 430 "?99 = 0".
adam@388 431 >>
adam@388 432
adam@398 433 The problem here is that variable [x] was introduced by [destruct] _after_ we introduced [?99] with [eexists], so the instantiation of [?99] may not mention [x]. A simple reordering of the proof solves the problem. *)
adam@388 434
adam@388 435 Restart.
adam@388 436 destruct 1 as [x]; apply ex_intro with x; symmetry; assumption.
adam@388 437 Qed.
adam@388 438
adam@429 439 (** This restriction for unification variables may seem counterintuitive, but it follows from the fact that CIC contains no concept of unification variable. Rather, to construct the final proof term, at the point in a proof where the unification variable is introduced, we replace it with the instantiation we eventually find for it. It is simply syntactically illegal to refer there to variables that are not in scope. Without such a restriction, we could trivially "prove" such non-theorems as [exists n : nat, forall m : nat, n = m] by [econstructor; intro; reflexivity]. *)
adam@388 440
adam@388 441
adamc@229 442 (** * The [Prop] Universe *)
adamc@229 443
adam@429 444 (** In Chapter 4, we saw parallel versions of useful datatypes for "programs" and "proofs." The convention was that programs live in [Set], and proofs live in [Prop]. We gave little explanation for why it is useful to maintain this distinction. There is certainly documentation value from separating programs from proofs; in practice, different concerns apply to building the two types of objects. It turns out, however, that these concerns motivate formal differences between the two universes in Coq.
adamc@229 445
adamc@229 446 Recall the types [sig] and [ex], which are the program and proof versions of existential quantification. Their definitions differ only in one place, where [sig] uses [Type] and [ex] uses [Prop]. *)
adamc@229 447
adamc@229 448 Print sig.
adamc@229 449 (** %\vspace{-.15in}% [[
adamc@229 450 Inductive sig (A : Type) (P : A -> Prop) : Type :=
adamc@229 451 exist : forall x : A, P x -> sig P
adam@302 452 ]]
adam@302 453 *)
adamc@229 454
adamc@229 455 Print ex.
adamc@229 456 (** %\vspace{-.15in}% [[
adamc@229 457 Inductive ex (A : Type) (P : A -> Prop) : Prop :=
adamc@229 458 ex_intro : forall x : A, P x -> ex P
adamc@229 459 ]]
adamc@229 460
adamc@229 461 It is natural to want a function to extract the first components of data structures like these. Doing so is easy enough for [sig]. *)
adamc@229 462
adamc@229 463 Definition projS A (P : A -> Prop) (x : sig P) : A :=
adamc@229 464 match x with
adamc@229 465 | exist v _ => v
adamc@229 466 end.
adamc@229 467
adam@429 468 (* begin hide *)
adam@437 469 (* begin thide *)
adam@429 470 Definition projE := O.
adam@437 471 (* end thide *)
adam@429 472 (* end hide *)
adam@429 473
adamc@229 474 (** We run into trouble with a version that has been changed to work with [ex].
adamc@229 475 [[
adamc@229 476 Definition projE A (P : A -> Prop) (x : ex P) : A :=
adamc@229 477 match x with
adamc@229 478 | ex_intro v _ => v
adamc@229 479 end.
adam@343 480 ]]
adamc@229 481
adam@343 482 <<
adamc@229 483 Error:
adamc@229 484 Incorrect elimination of "x" in the inductive type "ex":
adamc@229 485 the return type has sort "Type" while it should be "Prop".
adamc@229 486 Elimination of an inductive object of sort Prop
adamc@229 487 is not allowed on a predicate in sort Type
adamc@229 488 because proofs can be eliminated only to build proofs.
adam@343 489 >>
adamc@229 490
adam@429 491 In formal Coq parlance, %\index{elimination}%"elimination" means "pattern-matching." The typing rules of Gallina forbid us from pattern-matching on a discriminee whose type belongs to [Prop], whenever the result type of the [match] has a type besides [Prop]. This is a sort of "information flow" policy, where the type system ensures that the details of proofs can never have any effect on parts of a development that are not also marked as proofs.
adamc@229 492
adamc@229 493 This restriction matches informal practice. We think of programs and proofs as clearly separated, and, outside of constructive logic, the idea of computing with proofs is ill-formed. The distinction also has practical importance in Coq, where it affects the behavior of extraction.
adamc@229 494
adam@398 495 Recall that %\index{program extraction}%extraction is Coq's facility for translating Coq developments into programs in general-purpose programming languages like OCaml. Extraction _erases_ proofs and leaves programs intact. A simple example with [sig] and [ex] demonstrates the distinction. *)
adamc@229 496
adamc@229 497 Definition sym_sig (x : sig (fun n => n = 0)) : sig (fun n => 0 = n) :=
adamc@229 498 match x with
adamc@229 499 | exist n pf => exist _ n (sym_eq pf)
adamc@229 500 end.
adamc@229 501
adamc@229 502 Extraction sym_sig.
adamc@229 503 (** <<
adamc@229 504 (** val sym_sig : nat -> nat **)
adamc@229 505
adamc@229 506 let sym_sig x = x
adamc@229 507 >>
adamc@229 508
adamc@229 509 Since extraction erases proofs, the second components of [sig] values are elided, making [sig] a simple identity type family. The [sym_sig] operation is thus an identity function. *)
adamc@229 510
adamc@229 511 Definition sym_ex (x : ex (fun n => n = 0)) : ex (fun n => 0 = n) :=
adamc@229 512 match x with
adamc@229 513 | ex_intro n pf => ex_intro _ n (sym_eq pf)
adamc@229 514 end.
adamc@229 515
adamc@229 516 Extraction sym_ex.
adamc@229 517 (** <<
adamc@229 518 (** val sym_ex : __ **)
adamc@229 519
adamc@229 520 let sym_ex = __
adamc@229 521 >>
adamc@229 522
adam@435 523 In this example, the [ex] type itself is in [Prop], so whole [ex] packages are erased. Coq extracts every proposition as the (Coq-specific) type <<__>>, whose single constructor is <<__>>. Not only are proofs replaced by [__], but proof arguments to functions are also removed completely, as we see here.
adamc@229 524
adam@419 525 Extraction is very helpful as an optimization over programs that contain proofs. In languages like Haskell, advanced features make it possible to program with proofs, as a way of convincing the type checker to accept particular definitions. Unfortunately, when proofs are encoded as values in GADTs%~\cite{GADT}%, these proofs exist at runtime and consume resources. In contrast, with Coq, as long as all proofs are kept within [Prop], extraction is guaranteed to erase them.
adamc@229 526
adam@398 527 Many fans of the %\index{Curry-Howard correspondence}%Curry-Howard correspondence support the idea of _extracting programs from proofs_. In reality, few users of Coq and related tools do any such thing. Instead, extraction is better thought of as an optimization that reduces the runtime costs of expressive typing.
adamc@229 528
adamc@229 529 %\medskip%
adamc@229 530
adam@409 531 We have seen two of the differences between proofs and programs: proofs are subject to an elimination restriction and are elided by extraction. The remaining difference is that [Prop] is%\index{impredicativity}% _impredicative_, as this example shows. *)
adamc@229 532
adamc@229 533 Check forall P Q : Prop, P \/ Q -> Q \/ P.
adamc@229 534 (** %\vspace{-.15in}% [[
adamc@229 535 forall P Q : Prop, P \/ Q -> Q \/ P
adamc@229 536 : Prop
adamc@229 537 ]]
adamc@229 538
adamc@230 539 We see that it is possible to define a [Prop] that quantifies over other [Prop]s. This is fortunate, as we start wanting that ability even for such basic purposes as stating propositional tautologies. In the next section of this chapter, we will see some reasons why unrestricted impredicativity is undesirable. The impredicativity of [Prop] interacts crucially with the elimination restriction to avoid those pitfalls.
adamc@230 540
adamc@230 541 Impredicativity also allows us to implement a version of our earlier [exp] type that does not suffer from the weakness that we found. *)
adamc@230 542
adamc@230 543 Inductive expP : Type -> Prop :=
adamc@230 544 | ConstP : forall T, T -> expP T
adamc@230 545 | PairP : forall T1 T2, expP T1 -> expP T2 -> expP (T1 * T2)
adamc@230 546 | EqP : forall T, expP T -> expP T -> expP bool.
adamc@230 547
adamc@230 548 Check ConstP 0.
adamc@230 549 (** %\vspace{-.15in}% [[
adamc@230 550 ConstP 0
adamc@230 551 : expP nat
adam@302 552 ]]
adam@302 553 *)
adamc@230 554
adamc@230 555 Check PairP (ConstP 0) (ConstP tt).
adamc@230 556 (** %\vspace{-.15in}% [[
adamc@230 557 PairP (ConstP 0) (ConstP tt)
adamc@230 558 : expP (nat * unit)
adam@302 559 ]]
adam@302 560 *)
adamc@230 561
adamc@230 562 Check EqP (ConstP Set) (ConstP Type).
adamc@230 563 (** %\vspace{-.15in}% [[
adamc@230 564 EqP (ConstP Set) (ConstP Type)
adamc@230 565 : expP bool
adam@302 566 ]]
adam@302 567 *)
adamc@230 568
adamc@230 569 Check ConstP (ConstP O).
adamc@230 570 (** %\vspace{-.15in}% [[
adamc@230 571 ConstP (ConstP 0)
adamc@230 572 : expP (expP nat)
adamc@230 573 ]]
adamc@230 574
adam@287 575 In this case, our victory is really a shallow one. As we have marked [expP] as a family of proofs, we cannot deconstruct our expressions in the usual programmatic ways, which makes them almost useless for the usual purposes. Impredicative quantification is much more useful in defining inductive families that we really think of as judgments. For instance, this code defines a notion of equality that is strictly more permissive than the base equality [=]. *)
adamc@230 576
adamc@230 577 Inductive eqPlus : forall T, T -> T -> Prop :=
adamc@230 578 | Base : forall T (x : T), eqPlus x x
adamc@230 579 | Func : forall dom ran (f1 f2 : dom -> ran),
adamc@230 580 (forall x : dom, eqPlus (f1 x) (f2 x))
adamc@230 581 -> eqPlus f1 f2.
adamc@230 582
adamc@230 583 Check (Base 0).
adamc@230 584 (** %\vspace{-.15in}% [[
adamc@230 585 Base 0
adamc@230 586 : eqPlus 0 0
adam@302 587 ]]
adam@302 588 *)
adamc@230 589
adamc@230 590 Check (Func (fun n => n) (fun n => 0 + n) (fun n => Base n)).
adamc@230 591 (** %\vspace{-.15in}% [[
adamc@230 592 Func (fun n : nat => n) (fun n : nat => 0 + n) (fun n : nat => Base n)
adamc@230 593 : eqPlus (fun n : nat => n) (fun n : nat => 0 + n)
adam@302 594 ]]
adam@302 595 *)
adamc@230 596
adamc@230 597 Check (Base (Base 1)).
adamc@230 598 (** %\vspace{-.15in}% [[
adamc@230 599 Base (Base 1)
adamc@230 600 : eqPlus (Base 1) (Base 1)
adam@302 601 ]]
adam@302 602 *)
adamc@230 603
adam@343 604 (** Stating equality facts about proofs may seem baroque, but we have already seen its utility in the chapter on reasoning about equality proofs. *)
adam@343 605
adamc@230 606
adamc@230 607 (** * Axioms *)
adamc@230 608
adam@409 609 (** While the specific logic Gallina is hardcoded into Coq's implementation, it is possible to add certain logical rules in a controlled way. In other words, Coq may be used to reason about many different refinements of Gallina where strictly more theorems are provable. We achieve this by asserting%\index{axioms}% _axioms_ without proof.
adamc@230 610
adamc@230 611 We will motivate the idea by touring through some standard axioms, as enumerated in Coq's online FAQ. I will add additional commentary as appropriate. *)
adamc@230 612
adamc@230 613 (** ** The Basics *)
adamc@230 614
adam@343 615 (** One simple example of a useful axiom is the %\index{law of the excluded middle}%law of the excluded middle. *)
adamc@230 616
adamc@230 617 Require Import Classical_Prop.
adamc@230 618 Print classic.
adamc@230 619 (** %\vspace{-.15in}% [[
adamc@230 620 *** [ classic : forall P : Prop, P \/ ~ P ]
adamc@230 621 ]]
adamc@230 622
adam@343 623 In the implementation of module [Classical_Prop], this axiom was defined with the command%\index{Vernacular commands!Axiom}% *)
adamc@230 624
adamc@230 625 Axiom classic : forall P : Prop, P \/ ~ P.
adamc@230 626
adam@343 627 (** An [Axiom] may be declared with any type, in any of the universes. There is a synonym %\index{Vernacular commands!Parameter}%[Parameter] for [Axiom], and that synonym is often clearer for assertions not of type [Prop]. For instance, we can assert the existence of objects with certain properties. *)
adamc@230 628
adam@458 629 Parameter num : nat.
adam@458 630 Axiom positive : num > 0.
adam@458 631 Reset num.
adamc@230 632
adam@429 633 (** This kind of "axiomatic presentation" of a theory is very common outside of higher-order logic. However, in Coq, it is almost always preferable to stick to defining your objects, functions, and predicates via inductive definitions and functional programming.
adamc@230 634
adam@409 635 In general, there is a significant burden associated with any use of axioms. It is easy to assert a set of axioms that together is%\index{inconsistent axioms}% _inconsistent_. That is, a set of axioms may imply [False], which allows any theorem to be proved, which defeats the purpose of a proof assistant. For example, we could assert the following axiom, which is consistent by itself but inconsistent when combined with [classic]. *)
adamc@230 636
adam@287 637 Axiom not_classic : ~ forall P : Prop, P \/ ~ P.
adamc@230 638
adamc@230 639 Theorem uhoh : False.
adam@287 640 generalize classic not_classic; tauto.
adamc@230 641 Qed.
adamc@230 642
adamc@230 643 Theorem uhoh_again : 1 + 1 = 3.
adamc@230 644 destruct uhoh.
adamc@230 645 Qed.
adamc@230 646
adamc@230 647 Reset not_classic.
adamc@230 648
adam@429 649 (** On the subject of the law of the excluded middle itself, this axiom is usually quite harmless, and many practical Coq developments assume it. It has been proved metatheoretically to be consistent with CIC. Here, "proved metatheoretically" means that someone proved on paper that excluded middle holds in a _model_ of CIC in set theory%~\cite{SetsInTypes}%. All of the other axioms that we will survey in this section hold in the same model, so they are all consistent together.
adamc@230 650
adam@475 651 Recall that Coq implements%\index{constructive logic}% _constructive_ logic by default, where the law of the excluded middle is not provable. Proofs in constructive logic can be thought of as programs. A [forall] quantifier denotes a dependent function type, and a disjunction denotes a variant type. In such a setting, excluded middle could be interpreted as a decision procedure for arbitrary propositions, which computability theory tells us cannot exist. Thus, constructive logic with excluded middle can no longer be associated with our usual notion of programming.
adamc@230 652
adam@398 653 Given all this, why is it all right to assert excluded middle as an axiom? The intuitive justification is that the elimination restriction for [Prop] prevents us from treating proofs as programs. An excluded middle axiom that quantified over [Set] instead of [Prop] _would_ be problematic. If a development used that axiom, we would not be able to extract the code to OCaml (soundly) without implementing a genuine universal decision procedure. In contrast, values whose types belong to [Prop] are always erased by extraction, so we sidestep the axiom's algorithmic consequences.
adamc@230 654
adam@343 655 Because the proper use of axioms is so precarious, there are helpful commands for determining which axioms a theorem relies on.%\index{Vernacular commands!Print Assumptions}% *)
adamc@230 656
adamc@230 657 Theorem t1 : forall P : Prop, P -> ~ ~ P.
adamc@230 658 tauto.
adamc@230 659 Qed.
adamc@230 660
adamc@230 661 Print Assumptions t1.
adam@343 662 (** <<
adamc@230 663 Closed under the global context
adam@343 664 >>
adam@302 665 *)
adamc@230 666
adamc@230 667 Theorem t2 : forall P : Prop, ~ ~ P -> P.
adam@444 668 (** %\vspace{-.25in}%[[
adamc@230 669 tauto.
adam@343 670 ]]
adam@343 671 <<
adamc@230 672 Error: tauto failed.
adam@343 673 >>
adam@302 674 *)
adamc@230 675 intro P; destruct (classic P); tauto.
adamc@230 676 Qed.
adamc@230 677
adamc@230 678 Print Assumptions t2.
adamc@230 679 (** %\vspace{-.15in}% [[
adamc@230 680 Axioms:
adamc@230 681 classic : forall P : Prop, P \/ ~ P
adamc@230 682 ]]
adamc@230 683
adam@398 684 It is possible to avoid this dependence in some specific cases, where excluded middle _is_ provable, for decidable families of propositions. *)
adamc@230 685
adam@287 686 Theorem nat_eq_dec : forall n m : nat, n = m \/ n <> m.
adamc@230 687 induction n; destruct m; intuition; generalize (IHn m); intuition.
adamc@230 688 Qed.
adamc@230 689
adamc@230 690 Theorem t2' : forall n m : nat, ~ ~ (n = m) -> n = m.
adam@287 691 intros n m; destruct (nat_eq_dec n m); tauto.
adamc@230 692 Qed.
adamc@230 693
adamc@230 694 Print Assumptions t2'.
adam@343 695 (** <<
adamc@230 696 Closed under the global context
adam@343 697 >>
adamc@230 698
adamc@230 699 %\bigskip%
adamc@230 700
adam@409 701 Mainstream mathematical practice assumes excluded middle, so it can be useful to have it available in Coq developments, though it is also nice to know that a theorem is proved in a simpler formal system than classical logic. There is a similar story for%\index{proof irrelevance}% _proof irrelevance_, which simplifies proof issues that would not even arise in mainstream math. *)
adamc@230 702
adamc@230 703 Require Import ProofIrrelevance.
adamc@230 704 Print proof_irrelevance.
adam@458 705
adamc@230 706 (** %\vspace{-.15in}% [[
adamc@230 707 *** [ proof_irrelevance : forall (P : Prop) (p1 p2 : P), p1 = p2 ]
adamc@230 708 ]]
adamc@230 709
adam@458 710 This axiom asserts that any two proofs of the same proposition are equal. Recall this example function from Chapter 6. *)
adamc@230 711
adamc@230 712 (* begin hide *)
adamc@230 713 Lemma zgtz : 0 > 0 -> False.
adamc@230 714 crush.
adamc@230 715 Qed.
adamc@230 716 (* end hide *)
adamc@230 717
adamc@230 718 Definition pred_strong1 (n : nat) : n > 0 -> nat :=
adamc@230 719 match n with
adamc@230 720 | O => fun pf : 0 > 0 => match zgtz pf with end
adamc@230 721 | S n' => fun _ => n'
adamc@230 722 end.
adamc@230 723
adam@343 724 (** We might want to prove that different proofs of [n > 0] do not lead to different results from our richly typed predecessor function. *)
adamc@230 725
adamc@230 726 Theorem pred_strong1_irrel : forall n (pf1 pf2 : n > 0), pred_strong1 pf1 = pred_strong1 pf2.
adamc@230 727 destruct n; crush.
adamc@230 728 Qed.
adamc@230 729
adamc@230 730 (** The proof script is simple, but it involved peeking into the definition of [pred_strong1]. For more complicated function definitions, it can be considerably more work to prove that they do not discriminate on details of proof arguments. This can seem like a shame, since the [Prop] elimination restriction makes it impossible to write any function that does otherwise. Unfortunately, this fact is only true metatheoretically, unless we assert an axiom like [proof_irrelevance]. With that axiom, we can prove our theorem without consulting the definition of [pred_strong1]. *)
adamc@230 731
adamc@230 732 Theorem pred_strong1_irrel' : forall n (pf1 pf2 : n > 0), pred_strong1 pf1 = pred_strong1 pf2.
adamc@230 733 intros; f_equal; apply proof_irrelevance.
adamc@230 734 Qed.
adamc@230 735
adamc@230 736
adamc@230 737 (** %\bigskip%
adamc@230 738
adamc@230 739 In the chapter on equality, we already discussed some axioms that are related to proof irrelevance. In particular, Coq's standard library includes this axiom: *)
adamc@230 740
adamc@230 741 Require Import Eqdep.
adamc@230 742 Import Eq_rect_eq.
adamc@230 743 Print eq_rect_eq.
adamc@230 744 (** %\vspace{-.15in}% [[
adamc@230 745 *** [ eq_rect_eq :
adamc@230 746 forall (U : Type) (p : U) (Q : U -> Type) (x : Q p) (h : p = p),
adamc@230 747 x = eq_rect p Q x p h ]
adamc@230 748 ]]
adamc@230 749
adam@429 750 This axiom says that it is permissible to simplify pattern matches over proofs of equalities like [e = e]. The axiom is logically equivalent to some simpler corollaries. In the theorem names, "UIP" stands for %\index{unicity of identity proofs}%"unicity of identity proofs", where "identity" is a synonym for "equality." *)
adamc@230 751
adam@426 752 Corollary UIP_refl : forall A (x : A) (pf : x = x), pf = eq_refl x.
adam@426 753 intros; replace pf with (eq_rect x (eq x) (eq_refl x) x pf); [
adamc@230 754 symmetry; apply eq_rect_eq
adamc@230 755 | exact (match pf as pf' return match pf' in _ = y return x = y with
adam@426 756 | eq_refl => eq_refl x
adamc@230 757 end = pf' with
adam@426 758 | eq_refl => eq_refl _
adamc@230 759 end) ].
adamc@230 760 Qed.
adamc@230 761
adamc@230 762 Corollary UIP : forall A (x y : A) (pf1 pf2 : x = y), pf1 = pf2.
adamc@230 763 intros; generalize pf1 pf2; subst; intros;
adamc@230 764 match goal with
adamc@230 765 | [ |- ?pf1 = ?pf2 ] => rewrite (UIP_refl pf1); rewrite (UIP_refl pf2); reflexivity
adamc@230 766 end.
adamc@230 767 Qed.
adamc@230 768
adam@436 769 (* begin hide *)
adam@437 770 (* begin thide *)
adam@436 771 Require Eqdep_dec.
adam@437 772 (* end thide *)
adam@436 773 (* end hide *)
adam@436 774
adamc@231 775 (** These corollaries are special cases of proof irrelevance. In developments that only need proof irrelevance for equality, there is no need to assert full irrelevance.
adamc@230 776
adamc@230 777 Another facet of proof irrelevance is that, like excluded middle, it is often provable for specific propositions. For instance, [UIP] is provable whenever the type [A] has a decidable equality operation. The module [Eqdep_dec] of the standard library contains a proof. A similar phenomenon applies to other notable cases, including less-than proofs. Thus, it is often possible to use proof irrelevance without asserting axioms.
adamc@230 778
adamc@230 779 %\bigskip%
adamc@230 780
adamc@230 781 There are two more basic axioms that are often assumed, to avoid complications that do not arise in set theory. *)
adamc@230 782
adamc@230 783 Require Import FunctionalExtensionality.
adamc@230 784 Print functional_extensionality_dep.
adamc@230 785 (** %\vspace{-.15in}% [[
adamc@230 786 *** [ functional_extensionality_dep :
adamc@230 787 forall (A : Type) (B : A -> Type) (f g : forall x : A, B x),
adamc@230 788 (forall x : A, f x = g x) -> f = g ]
adamc@230 789
adamc@230 790 ]]
adamc@230 791
adamc@230 792 This axiom says that two functions are equal if they map equal inputs to equal outputs. Such facts are not provable in general in CIC, but it is consistent to assume that they are.
adamc@230 793
adam@343 794 A simple corollary shows that the same property applies to predicates. *)
adamc@230 795
adamc@230 796 Corollary predicate_extensionality : forall (A : Type) (B : A -> Prop) (f g : forall x : A, B x),
adamc@230 797 (forall x : A, f x = g x) -> f = g.
adamc@230 798 intros; apply functional_extensionality_dep; assumption.
adamc@230 799 Qed.
adamc@230 800
adam@343 801 (** In some cases, one might prefer to assert this corollary as the axiom, to restrict the consequences to proofs and not programs. *)
adam@343 802
adamc@230 803
adamc@230 804 (** ** Axioms of Choice *)
adamc@230 805
adam@343 806 (** Some Coq axioms are also points of contention in mainstream math. The most prominent example is the %\index{axiom of choice}%axiom of choice. In fact, there are multiple versions that we might consider, and, considered in isolation, none of these versions means quite what it means in classical set theory.
adamc@230 807
adam@398 808 First, it is possible to implement a choice operator _without_ axioms in some potentially surprising cases. *)
adamc@230 809
adamc@230 810 Require Import ConstructiveEpsilon.
adamc@230 811 Check constructive_definite_description.
adamc@230 812 (** %\vspace{-.15in}% [[
adamc@230 813 constructive_definite_description
adamc@230 814 : forall (A : Set) (f : A -> nat) (g : nat -> A),
adamc@230 815 (forall x : A, g (f x) = x) ->
adamc@230 816 forall P : A -> Prop,
adam@505 817 (forall x : A, {P x} + { ~ P x}) ->
adamc@230 818 (exists! x : A, P x) -> {x : A | P x}
adam@302 819 ]]
adam@302 820 *)
adamc@230 821
adamc@230 822 Print Assumptions constructive_definite_description.
adam@343 823 (** <<
adamc@230 824 Closed under the global context
adam@343 825 >>
adamc@230 826
adam@398 827 This function transforms a decidable predicate [P] into a function that produces an element satisfying [P] from a proof that such an element exists. The functions [f] and [g], in conjunction with an associated injectivity property, are used to express the idea that the set [A] is countable. Under these conditions, a simple brute force algorithm gets the job done: we just enumerate all elements of [A], stopping when we find one satisfying [P]. The existence proof, specified in terms of _unique_ existence [exists!], guarantees termination. The definition of this operator in Coq uses some interesting techniques, as seen in the implementation of the [ConstructiveEpsilon] module.
adamc@230 828
adamc@230 829 Countable choice is provable in set theory without appealing to the general axiom of choice. To support the more general principle in Coq, we must also add an axiom. Here is a functional version of the axiom of unique choice. *)
adamc@230 830
adamc@230 831 Require Import ClassicalUniqueChoice.
adamc@230 832 Check dependent_unique_choice.
adamc@230 833 (** %\vspace{-.15in}% [[
adamc@230 834 dependent_unique_choice
adamc@230 835 : forall (A : Type) (B : A -> Type) (R : forall x : A, B x -> Prop),
adamc@230 836 (forall x : A, exists! y : B x, R x y) ->
adam@343 837 exists f : forall x : A, B x,
adam@343 838 forall x : A, R x (f x)
adamc@230 839 ]]
adamc@230 840
adamc@230 841 This axiom lets us convert a relational specification [R] into a function implementing that specification. We need only prove that [R] is truly a function. An alternate, stronger formulation applies to cases where [R] maps each input to one or more outputs. We also simplify the statement of the theorem by considering only non-dependent function types. *)
adamc@230 842
adam@436 843 (* begin hide *)
adam@437 844 (* begin thide *)
adam@436 845 Require RelationalChoice.
adam@437 846 (* end thide *)
adam@436 847 (* end hide *)
adam@436 848
adamc@230 849 Require Import ClassicalChoice.
adamc@230 850 Check choice.
adamc@230 851 (** %\vspace{-.15in}% [[
adamc@230 852 choice
adamc@230 853 : forall (A B : Type) (R : A -> B -> Prop),
adamc@230 854 (forall x : A, exists y : B, R x y) ->
adamc@230 855 exists f : A -> B, forall x : A, R x (f x)
adam@444 856 ]]
adamc@230 857
adamc@230 858 This principle is proved as a theorem, based on the unique choice axiom and an additional axiom of relational choice from the [RelationalChoice] module.
adamc@230 859
adamc@230 860 In set theory, the axiom of choice is a fundamental philosophical commitment one makes about the universe of sets. In Coq, the choice axioms say something weaker. For instance, consider the simple restatement of the [choice] axiom where we replace existential quantification by its Curry-Howard analogue, subset types. *)
adamc@230 861
adamc@230 862 Definition choice_Set (A B : Type) (R : A -> B -> Prop) (H : forall x : A, {y : B | R x y})
adamc@230 863 : {f : A -> B | forall x : A, R x (f x)} :=
adamc@230 864 exist (fun f => forall x : A, R x (f x))
adamc@230 865 (fun x => proj1_sig (H x)) (fun x => proj2_sig (H x)).
adamc@230 866
adam@458 867 (** %\smallskip{}%Via the Curry-Howard correspondence, this "axiom" can be taken to have the same meaning as the original. It is implemented trivially as a transformation not much deeper than uncurrying. Thus, we see that the utility of the axioms that we mentioned earlier comes in their usage to build programs from proofs. Normal set theory has no explicit proofs, so the meaning of the usual axiom of choice is subtly different. In Gallina, the axioms implement a controlled relaxation of the restrictions on information flow from proofs to programs.
adamc@230 868
adam@505 869 However, when we combine an axiom of choice with the law of the excluded middle, the idea of "choice" becomes more interesting. Excluded middle gives us a highly non-computational way of constructing proofs, but it does not change the computational nature of programs. Thus, the axiom of choice is still giving us a way of translating between two different sorts of "programs," but the input programs (which are proofs) may be written in a rich language that goes beyond normal computability. This combination truly is more than repackaging a function with a different type.
adamc@230 870
adamc@230 871 %\bigskip%
adamc@230 872
adam@505 873 The Coq tools support a command-line flag %\index{impredicative Set}%<<-impredicative-set>>, which modifies Gallina in a more fundamental way by making [Set] impredicative. A term like [forall T : Set, T] has type [Set], and inductive definitions in [Set] may have constructors that quantify over arguments of any types. To maintain consistency, an elimination restriction must be imposed, similarly to the restriction for [Prop]. The restriction only applies to large inductive types, where some constructor quantifies over a type of type [Type]. In such cases, a value in this inductive type may only be pattern-matched over to yield a result type whose type is [Set] or [Prop]. This rule contrasts with the rule for [Prop], where the restriction applies even to non-large inductive types, and where the result type may only have type [Prop].
adamc@230 874
adam@505 875 In old versions of Coq, [Set] was impredicative by default. Later versions make [Set] predicative to avoid inconsistency with some classical axioms. In particular, one should watch out when using impredicative [Set] with axioms of choice. In combination with excluded middle or predicate extensionality, inconsistency can result. Impredicative [Set] can be useful for modeling inherently impredicative mathematical concepts, but almost all Coq developments get by fine without it. *)
adamc@230 876
adamc@230 877 (** ** Axioms and Computation *)
adamc@230 878
adam@398 879 (** One additional axiom-related wrinkle arises from an aspect of Gallina that is very different from set theory: a notion of _computational equivalence_ is central to the definition of the formal system. Axioms tend not to play well with computation. Consider this example. We start by implementing a function that uses a type equality proof to perform a safe type-cast. *)
adamc@230 880
adamc@230 881 Definition cast (x y : Set) (pf : x = y) (v : x) : y :=
adamc@230 882 match pf with
adam@426 883 | eq_refl => v
adamc@230 884 end.
adamc@230 885
adamc@230 886 (** Computation over programs that use [cast] can proceed smoothly. *)
adamc@230 887
adam@426 888 Eval compute in (cast (eq_refl (nat -> nat)) (fun n => S n)) 12.
adam@343 889 (** %\vspace{-.15in}%[[
adamc@230 890 = 13
adamc@230 891 : nat
adam@302 892 ]]
adam@302 893 *)
adamc@230 894
adamc@230 895 (** Things do not go as smoothly when we use [cast] with proofs that rely on axioms. *)
adamc@230 896
adamc@230 897 Theorem t3 : (forall n : nat, fin (S n)) = (forall n : nat, fin (n + 1)).
adamc@230 898 change ((forall n : nat, (fun n => fin (S n)) n) = (forall n : nat, (fun n => fin (n + 1)) n));
adamc@230 899 rewrite (functional_extensionality (fun n => fin (n + 1)) (fun n => fin (S n))); crush.
adamc@230 900 Qed.
adamc@230 901
adamc@230 902 Eval compute in (cast t3 (fun _ => First)) 12.
adam@444 903 (** %\vspace{-.15in}%[[
adamc@230 904 = match t3 in (_ = P) return P with
adam@426 905 | eq_refl => fun n : nat => First
adamc@230 906 end 12
adamc@230 907 : fin (12 + 1)
adamc@230 908 ]]
adamc@230 909
adam@458 910 Computation gets stuck in a pattern-match on the proof [t3]. The structure of [t3] is not known, so the match cannot proceed. It turns out a more basic problem leads to this particular situation. We ended the proof of [t3] with [Qed], so the definition of [t3] is not available to computation. That mistake is easily fixed. *)
adamc@230 911
adamc@230 912 Reset t3.
adamc@230 913
adamc@230 914 Theorem t3 : (forall n : nat, fin (S n)) = (forall n : nat, fin (n + 1)).
adamc@230 915 change ((forall n : nat, (fun n => fin (S n)) n) = (forall n : nat, (fun n => fin (n + 1)) n));
adamc@230 916 rewrite (functional_extensionality (fun n => fin (n + 1)) (fun n => fin (S n))); crush.
adamc@230 917 Defined.
adamc@230 918
adamc@230 919 Eval compute in (cast t3 (fun _ => First)) 12.
adam@444 920 (** %\vspace{-.15in}%[[
adamc@230 921 = match
adamc@230 922 match
adamc@230 923 match
adamc@230 924 functional_extensionality
adamc@230 925 ....
adamc@230 926 ]]
adamc@230 927
adam@398 928 We elide most of the details. A very unwieldy tree of nested matches on equality proofs appears. This time evaluation really _is_ stuck on a use of an axiom.
adamc@230 929
adamc@230 930 If we are careful in using tactics to prove an equality, we can still compute with casts over the proof. *)
adamc@230 931
adamc@230 932 Lemma plus1 : forall n, S n = n + 1.
adamc@230 933 induction n; simpl; intuition.
adamc@230 934 Defined.
adamc@230 935
adamc@230 936 Theorem t4 : forall n, fin (S n) = fin (n + 1).
adamc@230 937 intro; f_equal; apply plus1.
adamc@230 938 Defined.
adamc@230 939
adamc@230 940 Eval compute in cast (t4 13) First.
adamc@230 941 (** %\vspace{-.15in}% [[
adamc@230 942 = First
adamc@230 943 : fin (13 + 1)
adam@302 944 ]]
adam@343 945
adam@426 946 This simple computational reduction hides the use of a recursive function to produce a suitable [eq_refl] proof term. The recursion originates in our use of [induction] in [t4]'s proof. *)
adam@343 947
adam@344 948
adam@344 949 (** ** Methods for Avoiding Axioms *)
adam@344 950
adam@409 951 (** The last section demonstrated one reason to avoid axioms: they interfere with computational behavior of terms. A further reason is to reduce the philosophical commitment of a theorem. The more axioms one assumes, the harder it becomes to convince oneself that the formal system corresponds appropriately to one's intuitions. A refinement of this last point, in applications like %\index{proof-carrying code}%proof-carrying code%~\cite{PCC}% in computer security, has to do with minimizing the size of a%\index{trusted code base}% _trusted code base_. To convince ourselves that a theorem is true, we must convince ourselves of the correctness of the program that checks the theorem. Axioms effectively become new source code for the checking program, increasing the effort required to perform a correctness audit.
adam@344 952
adam@429 953 An earlier section gave one example of avoiding an axiom. We proved that [pred_strong1] is agnostic to details of the proofs passed to it as arguments, by unfolding the definition of the function. A "simpler" proof keeps the function definition opaque and instead applies a proof irrelevance axiom. By accepting a more complex proof, we reduce our philosophical commitment and trusted base. (By the way, the less-than relation that the proofs in question here prove turns out to admit proof irrelevance as a theorem provable within normal Gallina!)
adam@344 954
adam@344 955 One dark secret of the [dep_destruct] tactic that we have used several times is reliance on an axiom. Consider this simple case analysis principle for [fin] values: *)
adam@344 956
adam@344 957 Theorem fin_cases : forall n (f : fin (S n)), f = First \/ exists f', f = Next f'.
adam@344 958 intros; dep_destruct f; eauto.
adam@344 959 Qed.
adam@344 960
adam@429 961 (* begin hide *)
adam@429 962 Require Import JMeq.
adam@437 963 (* begin thide *)
adam@429 964 Definition jme := (JMeq, JMeq_eq).
adam@437 965 (* end thide *)
adam@429 966 (* end hide *)
adam@429 967
adam@344 968 Print Assumptions fin_cases.
adam@344 969 (** %\vspace{-.15in}%[[
adam@344 970 Axioms:
adam@429 971 JMeq_eq : forall (A : Type) (x y : A), JMeq x y -> x = y
adam@344 972 ]]
adam@344 973
adam@344 974 The proof depends on the [JMeq_eq] axiom that we met in the chapter on equality proofs. However, a smarter tactic could have avoided an axiom dependence. Here is an alternate proof via a slightly strange looking lemma. *)
adam@344 975
adam@344 976 (* begin thide *)
adam@344 977 Lemma fin_cases_again' : forall n (f : fin n),
adam@344 978 match n return fin n -> Prop with
adam@344 979 | O => fun _ => False
adam@344 980 | S n' => fun f => f = First \/ exists f', f = Next f'
adam@344 981 end f.
adam@344 982 destruct f; eauto.
adam@344 983 Qed.
adam@344 984
adam@344 985 (** We apply a variant of the %\index{convoy pattern}%convoy pattern, which we are used to seeing in function implementations. Here, the pattern helps us state a lemma in a form where the argument to [fin] is a variable. Recall that, thanks to basic typing rules for pattern-matching, [destruct] will only work effectively on types whose non-parameter arguments are variables. The %\index{tactics!exact}%[exact] tactic, which takes as argument a literal proof term, now gives us an easy way of proving the original theorem. *)
adam@344 986
adam@344 987 Theorem fin_cases_again : forall n (f : fin (S n)), f = First \/ exists f', f = Next f'.
adam@344 988 intros; exact (fin_cases_again' f).
adam@344 989 Qed.
adam@344 990 (* end thide *)
adam@344 991
adam@344 992 Print Assumptions fin_cases_again.
adam@344 993 (** %\vspace{-.15in}%
adam@344 994 <<
adam@344 995 Closed under the global context
adam@344 996 >>
adam@344 997
adam@345 998 *)
adam@345 999
adam@345 1000 (* begin thide *)
adam@345 1001 (** As the Curry-Howard correspondence might lead us to expect, the same pattern may be applied in programming as in proving. Axioms are relevant in programming, too, because, while Coq includes useful extensions like [Program] that make dependently typed programming more straightforward, in general these extensions generate code that relies on axioms about equality. We can use clever pattern matching to write our code axiom-free.
adam@345 1002
adam@429 1003 As an example, consider a [Set] version of [fin_cases]. We use [Set] types instead of [Prop] types, so that return values have computational content and may be used to guide the behavior of algorithms. Beside that, we are essentially writing the same "proof" in a more explicit way. *)
adam@345 1004
adam@345 1005 Definition finOut n (f : fin n) : match n return fin n -> Type with
adam@345 1006 | O => fun _ => Empty_set
adam@345 1007 | _ => fun f => {f' : _ | f = Next f'} + {f = First}
adam@345 1008 end f :=
adam@345 1009 match f with
adam@426 1010 | First _ => inright _ (eq_refl _)
adam@426 1011 | Next _ f' => inleft _ (exist _ f' (eq_refl _))
adam@345 1012 end.
adam@345 1013 (* end thide *)
adam@345 1014
adam@345 1015 (** As another example, consider the following type of formulas in first-order logic. The intent of the type definition will not be important in what follows, but we give a quick intuition for the curious reader. Our formulas may include [forall] quantification over arbitrary [Type]s, and we index formulas by environments telling which variables are in scope and what their types are; such an environment is a [list Type]. A constructor [Inject] lets us include any Coq [Prop] as a formula, and [VarEq] and [Lift] can be used for variable references, in what is essentially the de Bruijn index convention. (Again, the detail in this paragraph is not important to understand the discussion that follows!) *)
adam@344 1016
adam@344 1017 Inductive formula : list Type -> Type :=
adam@344 1018 | Inject : forall Ts, Prop -> formula Ts
adam@344 1019 | VarEq : forall T Ts, T -> formula (T :: Ts)
adam@344 1020 | Lift : forall T Ts, formula Ts -> formula (T :: Ts)
adam@344 1021 | Forall : forall T Ts, formula (T :: Ts) -> formula Ts
adam@344 1022 | And : forall Ts, formula Ts -> formula Ts -> formula Ts.
adam@344 1023
adam@344 1024 (** This example is based on my own experiences implementing variants of a program logic called XCAP%~\cite{XCAP}%, which also includes an inductive predicate for characterizing which formulas are provable. Here I include a pared-down version of such a predicate, with only two constructors, which is sufficient to illustrate certain tricky issues. *)
adam@344 1025
adam@344 1026 Inductive proof : formula nil -> Prop :=
adam@344 1027 | PInject : forall (P : Prop), P -> proof (Inject nil P)
adam@344 1028 | PAnd : forall p q, proof p -> proof q -> proof (And p q).
adam@344 1029
adam@429 1030 (** Let us prove a lemma showing that a "[P /\ Q -> P]" rule is derivable within the rules of [proof]. *)
adam@344 1031
adam@344 1032 Theorem proj1 : forall p q, proof (And p q) -> proof p.
adam@344 1033 destruct 1.
adam@344 1034 (** %\vspace{-.15in}%[[
adam@344 1035 p : formula nil
adam@344 1036 q : formula nil
adam@344 1037 P : Prop
adam@344 1038 H : P
adam@344 1039 ============================
adam@344 1040 proof p
adam@344 1041 ]]
adam@344 1042 *)
adam@344 1043
adam@344 1044 (** We are reminded that [induction] and [destruct] do not work effectively on types with non-variable arguments. The first subgoal, shown above, is clearly unprovable. (Consider the case where [p = Inject nil False].)
adam@344 1045
adam@344 1046 An application of the %\index{tactics!dependent destruction}%[dependent destruction] tactic (the basis for [dep_destruct]) solves the problem handily. We use a shorthand with the %\index{tactics!intros}%[intros] tactic that lets us use question marks for variable names that do not matter. *)
adam@344 1047
adam@344 1048 Restart.
adam@344 1049 Require Import Program.
adam@344 1050 intros ? ? H; dependent destruction H; auto.
adam@344 1051 Qed.
adam@344 1052
adam@344 1053 Print Assumptions proj1.
adam@344 1054 (** %\vspace{-.15in}%[[
adam@344 1055 Axioms:
adam@344 1056 eq_rect_eq : forall (U : Type) (p : U) (Q : U -> Type) (x : Q p) (h : p = p),
adam@344 1057 x = eq_rect p Q x p h
adam@344 1058 ]]
adam@344 1059
adam@344 1060 Unfortunately, that built-in tactic appeals to an axiom. It is still possible to avoid axioms by giving the proof via another odd-looking lemma. Here is a first attempt that fails at remaining axiom-free, using a common equality-based trick for supporting induction on non-variable arguments to type families. The trick works fine without axioms for datatypes more traditional than [formula], but we run into trouble with our current type. *)
adam@344 1061
adam@344 1062 Lemma proj1_again' : forall r, proof r
adam@344 1063 -> forall p q, r = And p q -> proof p.
adam@344 1064 destruct 1; crush.
adam@344 1065 (** %\vspace{-.15in}%[[
adam@344 1066 H0 : Inject [] P = And p q
adam@344 1067 ============================
adam@344 1068 proof p
adam@344 1069 ]]
adam@344 1070
adam@344 1071 The first goal looks reasonable. Hypothesis [H0] is clearly contradictory, as [discriminate] can show. *)
adam@344 1072
adam@547 1073 try discriminate. (* Note: Coq 8.6 is now solving this subgoal automatically!
adam@547 1074 * This line left here to keep everything working in
adam@563 1075 * 8.4 and 8.5. *)
adam@344 1076 (** %\vspace{-.15in}%[[
adam@344 1077 H : proof p
adam@344 1078 H1 : And p q = And p0 q0
adam@344 1079 ============================
adam@344 1080 proof p0
adam@344 1081 ]]
adam@344 1082
adam@344 1083 It looks like we are almost done. Hypothesis [H1] gives [p = p0] by injectivity of constructors, and then [H] finishes the case. *)
adam@344 1084
adam@344 1085 injection H1; intros.
adam@344 1086
adam@429 1087 (* begin hide *)
adam@437 1088 (* begin thide *)
adam@429 1089 Definition existT' := existT.
adam@437 1090 (* end thide *)
adam@429 1091 (* end hide *)
adam@429 1092
adam@429 1093 (** Unfortunately, the "equality" that we expected between [p] and [p0] comes in a strange form:
adam@344 1094
adam@344 1095 [[
adam@344 1096 H3 : existT (fun Ts : list Type => formula Ts) []%list p =
adam@344 1097 existT (fun Ts : list Type => formula Ts) []%list p0
adam@344 1098 ============================
adam@344 1099 proof p0
adam@344 1100 ]]
adam@344 1101
adam@345 1102 It may take a bit of tinkering, but, reviewing Chapter 3's discussion of writing injection principles manually, it makes sense that an [existT] type is the most direct way to express the output of [injection] on a dependently typed constructor. The constructor [And] is dependently typed, since it takes a parameter [Ts] upon which the types of [p] and [q] depend. Let us not dwell further here on why this goal appears; the reader may like to attempt the (impossible) exercise of building a better injection lemma for [And], without using axioms.
adam@344 1103
adam@344 1104 How exactly does an axiom come into the picture here? Let us ask [crush] to finish the proof. *)
adam@344 1105
adam@344 1106 crush.
adam@344 1107 Qed.
adam@344 1108
adam@344 1109 Print Assumptions proj1_again'.
adam@344 1110 (** %\vspace{-.15in}%[[
adam@344 1111 Axioms:
adam@344 1112 eq_rect_eq : forall (U : Type) (p : U) (Q : U -> Type) (x : Q p) (h : p = p),
adam@344 1113 x = eq_rect p Q x p h
adam@344 1114 ]]
adam@344 1115
adam@344 1116 It turns out that this familiar axiom about equality (or some other axiom) is required to deduce [p = p0] from the hypothesis [H3] above. The soundness of that proof step is neither provable nor disprovable in Gallina.
adam@344 1117
adam@479 1118 Hope is not lost, however. We can produce an even stranger looking lemma, which gives us the theorem without axioms. As always when we want to do case analysis on a term with a tricky dependent type, the key is to refactor the theorem statement so that every term we [match] on has _variables_ as its type indices; so instead of talking about proofs of [And p q], we talk about proofs of an arbitrary [r], but we only conclude anything interesting when [r] is an [And]. *)
adam@344 1119
adam@344 1120 Lemma proj1_again'' : forall r, proof r
adam@344 1121 -> match r with
adam@344 1122 | And Ps p _ => match Ps return formula Ps -> Prop with
adam@344 1123 | nil => fun p => proof p
adam@344 1124 | _ => fun _ => True
adam@344 1125 end p
adam@344 1126 | _ => True
adam@344 1127 end.
adam@344 1128 destruct 1; auto.
adam@344 1129 Qed.
adam@344 1130
adam@344 1131 Theorem proj1_again : forall p q, proof (And p q) -> proof p.
adam@344 1132 intros ? ? H; exact (proj1_again'' H).
adam@344 1133 Qed.
adam@344 1134
adam@344 1135 Print Assumptions proj1_again.
adam@344 1136 (** <<
adam@344 1137 Closed under the global context
adam@344 1138 >>
adam@344 1139
adam@377 1140 This example illustrates again how some of the same design patterns we learned for dependently typed programming can be used fruitfully in theorem statements.
adam@377 1141
adam@377 1142 %\medskip%
adam@377 1143
adam@398 1144 To close the chapter, we consider one final way to avoid dependence on axioms. Often this task is equivalent to writing definitions such that they _compute_. That is, we want Coq's normal reduction to be able to run certain programs to completion. Here is a simple example where such computation can get stuck. In proving properties of such functions, we would need to apply axioms like %\index{axiom K}%K manually to make progress.
adam@377 1145
adam@377 1146 Imagine we are working with %\index{deep embedding}%deeply embedded syntax of some programming language, where each term is considered to be in the scope of a number of free variables that hold normal Coq values. To enforce proper typing, we will need to model a Coq typing environment somehow. One natural choice is as a list of types, where variable number [i] will be treated as a reference to the [i]th element of the list. *)
adam@377 1147
adam@377 1148 Section withTypes.
adam@377 1149 Variable types : list Set.
adam@377 1150
adam@377 1151 (** To give the semantics of terms, we will need to represent value environments, which assign each variable a term of the proper type. *)
adam@377 1152
adam@377 1153 Variable values : hlist (fun x : Set => x) types.
adam@377 1154
adam@377 1155 (** Now imagine that we are writing some procedure that operates on a distinguished variable of type [nat]. A hypothesis formalizes this assumption, using the standard library function [nth_error] for looking up list elements by position. *)
adam@377 1156
adam@377 1157 Variable natIndex : nat.
adam@377 1158 Variable natIndex_ok : nth_error types natIndex = Some nat.
adam@377 1159
adam@377 1160 (** It is not hard to use this hypothesis to write a function for extracting the [nat] value in position [natIndex] of [values], starting with two helpful lemmas, each of which we finish with [Defined] to mark the lemma as transparent, so that its definition may be expanded during evaluation. *)
adam@377 1161
adam@377 1162 Lemma nth_error_nil : forall A n x,
adam@377 1163 nth_error (@nil A) n = Some x
adam@377 1164 -> False.
adam@377 1165 destruct n; simpl; unfold error; congruence.
adam@377 1166 Defined.
adam@377 1167
adam@377 1168 Implicit Arguments nth_error_nil [A n x].
adam@377 1169
adam@377 1170 Lemma Some_inj : forall A (x y : A),
adam@377 1171 Some x = Some y
adam@377 1172 -> x = y.
adam@377 1173 congruence.
adam@377 1174 Defined.
adam@377 1175
adam@377 1176 Fixpoint getNat (types' : list Set) (values' : hlist (fun x : Set => x) types')
adam@377 1177 (natIndex : nat) : (nth_error types' natIndex = Some nat) -> nat :=
adam@377 1178 match values' with
adam@377 1179 | HNil => fun pf => match nth_error_nil pf with end
adam@377 1180 | HCons t ts x values'' =>
adam@377 1181 match natIndex return nth_error (t :: ts) natIndex = Some nat -> nat with
adam@377 1182 | O => fun pf =>
adam@377 1183 match Some_inj pf in _ = T return T with
adam@426 1184 | eq_refl => x
adam@377 1185 end
adam@377 1186 | S natIndex' => getNat values'' natIndex'
adam@377 1187 end
adam@377 1188 end.
adam@377 1189 End withTypes.
adam@377 1190
adam@377 1191 (** The problem becomes apparent when we experiment with running [getNat] on a concrete [types] list. *)
adam@377 1192
adam@377 1193 Definition myTypes := unit :: nat :: bool :: nil.
adam@377 1194 Definition myValues : hlist (fun x : Set => x) myTypes :=
adam@377 1195 tt ::: 3 ::: false ::: HNil.
adam@377 1196
adam@377 1197 Definition myNatIndex := 1.
adam@377 1198
adam@377 1199 Theorem myNatIndex_ok : nth_error myTypes myNatIndex = Some nat.
adam@377 1200 reflexivity.
adam@377 1201 Defined.
adam@377 1202
adam@377 1203 Eval compute in getNat myValues myNatIndex myNatIndex_ok.
adam@377 1204 (** %\vspace{-.15in}%[[
adam@377 1205 = 3
adam@377 1206 ]]
adam@377 1207
adam@398 1208 We have not hit the problem yet, since we proceeded with a concrete equality proof for [myNatIndex_ok]. However, consider a case where we want to reason about the behavior of [getNat] _independently_ of a specific proof. *)
adam@377 1209
adam@377 1210 Theorem getNat_is_reasonable : forall pf, getNat myValues myNatIndex pf = 3.
adam@377 1211 intro; compute.
adam@377 1212 (**
adam@377 1213 <<
adam@377 1214 1 subgoal
adam@377 1215 >>
adam@377 1216 %\vspace{-.3in}%[[
adam@377 1217 pf : nth_error myTypes myNatIndex = Some nat
adam@377 1218 ============================
adam@377 1219 match
adam@377 1220 match
adam@377 1221 pf in (_ = y)
adam@377 1222 return (nat = match y with
adam@377 1223 | Some H => H
adam@377 1224 | None => nat
adam@377 1225 end)
adam@377 1226 with
adam@377 1227 | eq_refl => eq_refl
adam@377 1228 end in (_ = T) return T
adam@377 1229 with
adam@377 1230 | eq_refl => 3
adam@377 1231 end = 3
adam@377 1232 ]]
adam@377 1233
adam@377 1234 Since the details of the equality proof [pf] are not known, computation can proceed no further. A rewrite with axiom K would allow us to make progress, but we can rethink the definitions a bit to avoid depending on axioms. *)
adam@377 1235
adam@377 1236 Abort.
adam@377 1237
adam@377 1238 (** Here is a definition of a function that turns out to be useful, though no doubt its purpose will be mysterious for now. A call [update ls n x] overwrites the [n]th position of the list [ls] with the value [x], padding the end of the list with extra [x] values as needed to ensure sufficient length. *)
adam@377 1239
adam@377 1240 Fixpoint copies A (x : A) (n : nat) : list A :=
adam@377 1241 match n with
adam@377 1242 | O => nil
adam@377 1243 | S n' => x :: copies x n'
adam@377 1244 end.
adam@377 1245
adam@377 1246 Fixpoint update A (ls : list A) (n : nat) (x : A) : list A :=
adam@377 1247 match ls with
adam@377 1248 | nil => copies x n ++ x :: nil
adam@377 1249 | y :: ls' => match n with
adam@377 1250 | O => x :: ls'
adam@377 1251 | S n' => y :: update ls' n' x
adam@377 1252 end
adam@377 1253 end.
adam@377 1254
adam@377 1255 (** Now let us revisit the definition of [getNat]. *)
adam@377 1256
adam@377 1257 Section withTypes'.
adam@377 1258 Variable types : list Set.
adam@377 1259 Variable natIndex : nat.
adam@377 1260
adam@429 1261 (** Here is the trick: instead of asserting properties about the list [types], we build a "new" list that is _guaranteed by construction_ to have those properties. *)
adam@377 1262
adam@377 1263 Definition types' := update types natIndex nat.
adam@377 1264
adam@377 1265 Variable values : hlist (fun x : Set => x) types'.
adam@377 1266
adam@377 1267 (** Now a bit of dependent pattern matching helps us rewrite [getNat] in a way that avoids any use of equality proofs. *)
adam@377 1268
adam@378 1269 Fixpoint skipCopies (n : nat)
adam@378 1270 : hlist (fun x : Set => x) (copies nat n ++ nat :: nil) -> nat :=
adam@378 1271 match n with
adam@378 1272 | O => fun vs => hhd vs
adam@378 1273 | S n' => fun vs => skipCopies n' (htl vs)
adam@378 1274 end.
adam@378 1275
adam@377 1276 Fixpoint getNat' (types'' : list Set) (natIndex : nat)
adam@377 1277 : hlist (fun x : Set => x) (update types'' natIndex nat) -> nat :=
adam@377 1278 match types'' with
adam@378 1279 | nil => skipCopies natIndex
adam@377 1280 | t :: types0 =>
adam@377 1281 match natIndex return hlist (fun x : Set => x)
adam@377 1282 (update (t :: types0) natIndex nat) -> nat with
adam@377 1283 | O => fun vs => hhd vs
adam@377 1284 | S natIndex' => fun vs => getNat' types0 natIndex' (htl vs)
adam@377 1285 end
adam@377 1286 end.
adam@377 1287 End withTypes'.
adam@377 1288
adam@398 1289 (** Now the surprise comes in how easy it is to _use_ [getNat']. While typing works by modification of a types list, we can choose parameters so that the modification has no effect. *)
adam@377 1290
adam@377 1291 Theorem getNat_is_reasonable : getNat' myTypes myNatIndex myValues = 3.
adam@377 1292 reflexivity.
adam@377 1293 Qed.
adam@377 1294
adam@377 1295 (** The same parameters as before work without alteration, and we avoid use of axioms. *)