Mercurial > cpdt > repo

annotate src/MoreDep.v @ 380:31fa03bc0f18

Find changesets by keywords (author, files, the commit message), revision number or hash, or revset expression.
Get it to build with Coq 8.4
author Adam Chlipala <adam@chlipala.net>
date Thu, 29 Mar 2012 18:10:52 -0400
parents d1276004eec9
children b911d0df5eee
rev   line source
adam@371 1 (* Copyright (c) 2008-2012, Adam Chlipala
adamc@83 2 *
adamc@83 3 * This work is licensed under a
adamc@83 4 * Creative Commons Attribution-Noncommercial-No Derivative Works 3.0
adamc@83 5 * Unported License.
adamc@83 6 * The license text is available at:
adamc@83 7 * http://creativecommons.org/licenses/by-nc-nd/3.0/
adamc@83 8 *)
adamc@83 9
adamc@83 10 (* begin hide *)
adamc@85 11 Require Import Arith Bool List.
adamc@83 12
adam@314 13 Require Import CpdtTactics MoreSpecif.
adamc@83 14
adamc@83 15 Set Implicit Arguments.
adamc@83 16 (* end hide *)
adamc@83 17
adamc@83 18
adamc@83 19 (** %\chapter{More Dependent Types}% *)
adamc@83 20
adam@292 21 (** Subset types and their relatives help us integrate verification with programming. Though they reorganize the certified programmer's workflow, they tend not to have deep effects on proofs. We write largely the same proofs as we would for classical verification, with some of the structure moved into the programs themselves. It turns out that, when we use dependent types to their full potential, we warp the development and proving process even more than that, picking up %``%#"#free theorems#"#%''% to the extent that often a certified program is hardly more complex than its uncertified counterpart in Haskell or ML.
adamc@83 22
adam@338 23 In particular, we have only scratched the tip of the iceberg that is Coq's inductive definition mechanism. The inductive types we have seen so far have their counterparts in the other proof assistants that we surveyed in Chapter 1. This chapter explores the strange new world of dependent inductive datatypes (that is, dependent inductive types outside [Prop]), a possibility that sets Coq apart from all of the competition not based on type theory. *)
adamc@83 24
adamc@84 25
adamc@84 26 (** * Length-Indexed Lists *)
adamc@84 27
adam@338 28 (** Many introductions to dependent types start out by showing how to use them to eliminate array bounds checks%\index{array bounds checks}%. When the type of an array tells you how many elements it has, your compiler can detect out-of-bounds dereferences statically. Since we are working in a pure functional language, the next best thing is length-indexed lists%\index{length-indexed lists}%, which the following code defines. *)
adamc@84 29
adamc@84 30 Section ilist.
adamc@84 31 Variable A : Set.
adamc@84 32
adamc@84 33 Inductive ilist : nat -> Set :=
adamc@84 34 | Nil : ilist O
adamc@84 35 | Cons : forall n, A -> ilist n -> ilist (S n).
adamc@84 36
adamc@84 37 (** We see that, within its section, [ilist] is given type [nat -> Set]. Previously, every inductive type we have seen has either had plain [Set] as its type or has been a predicate with some type ending in [Prop]. The full generality of inductive definitions lets us integrate the expressivity of predicates directly into our normal programming.
adamc@84 38
adam@338 39 The [nat] argument to [ilist] tells us the length of the list. The types of [ilist]'s constructors tell us that a [Nil] list has length [O] and that a [Cons] list has length one greater than the length of its tail. We may apply [ilist] to any natural number, even natural numbers that are only known at runtime. It is this breaking of the %\index{phase distinction}\textit{%#<i>#phase distinction#</i>#%}% that characterizes [ilist] as %\textit{%#<i>#dependently typed#</i>#%}%.
adamc@84 40
adamc@213 41 In expositions of list types, we usually see the length function defined first, but here that would not be a very productive function to code. Instead, let us implement list concatenation. *)
adamc@84 42
adamc@213 43 Fixpoint app n1 (ls1 : ilist n1) n2 (ls2 : ilist n2) : ilist (n1 + n2) :=
adamc@213 44 match ls1 with
adamc@213 45 | Nil => ls2
adamc@213 46 | Cons _ x ls1' => Cons x (app ls1' ls2)
adamc@213 47 end.
adamc@84 48
adam@338 49 (** Past Coq versions signalled an error for this definition. The code is still invalid within Coq's core language, but current Coq versions automatically add annotations to the original program, producing a valid core program. These are the annotations on [match] discriminees that we began to study in the previous chapter. We can rewrite [app] to give the annotations explicitly. *)
adamc@100 50
adamc@100 51 (* begin thide *)
adam@338 52 Fixpoint app' n1 (ls1 : ilist n1) n2 (ls2 : ilist n2) : ilist (n1 + n2) :=
adam@338 53 match ls1 in (ilist n1) return (ilist (n1 + n2)) with
adam@338 54 | Nil => ls2
adam@338 55 | Cons _ x ls1' => Cons x (app' ls1' ls2)
adam@338 56 end.
adamc@100 57 (* end thide *)
adamc@84 58
adam@338 59 (** Using [return] alone allowed us to express a dependency of the [match] result type on the %\textit{%#<i>#value#</i>#%}% of the discriminee. What %\index{Gallina terms!in}%[in] adds to our arsenal is a way of expressing a dependency on the %\textit{%#<i>#type#</i>#%}% of the discriminee. Specifically, the [n1] in the [in] clause above is a %\textit{%#<i>#binding occurrence#</i>#%}% whose scope is the [return] clause.
adamc@84 60
adamc@84 61 We may use [in] clauses only to bind names for the arguments of an inductive type family. That is, each [in] clause must be an inductive type family name applied to a sequence of underscores and variable names of the proper length. The positions for %\textit{%#<i>#parameters#</i>#%}% to the type family must all be underscores. Parameters are those arguments declared with section variables or with entries to the left of the first colon in an inductive definition. They cannot vary depending on which constructor was used to build the discriminee, so Coq prohibits pointless matches on them. It is those arguments defined in the type to the right of the colon that we may name with [in] clauses.
adamc@84 62
adam@338 63 Our [app] function could be typed in so-called %\index{stratified type systems}\textit{%#<i>#stratified#</i>#%}% type systems, which avoid true dependency. That is, we could consider the length indices to lists to live in a separate, compile-time-only universe from the lists themselves. This stratification between a compile-time universe and a run-time universe, with no references to the latter in the former, gives rise to the terminology %``%#"#stratified.#"#%''% Our next example would be harder to implement in a stratified system. We write an injection function from regular lists to length-indexed lists. A stratified implementation would need to duplicate the definition of lists across compile-time and run-time versions, and the run-time versions would need to be indexed by the compile-time versions. *)
adamc@84 64
adamc@100 65 (* EX: Implement injection from normal lists *)
adamc@100 66
adamc@100 67 (* begin thide *)
adamc@84 68 Fixpoint inject (ls : list A) : ilist (length ls) :=
adamc@213 69 match ls with
adamc@84 70 | nil => Nil
adamc@84 71 | h :: t => Cons h (inject t)
adamc@84 72 end.
adamc@84 73
adamc@84 74 (** We can define an inverse conversion and prove that it really is an inverse. *)
adamc@84 75
adamc@213 76 Fixpoint unject n (ls : ilist n) : list A :=
adamc@84 77 match ls with
adamc@84 78 | Nil => nil
adamc@84 79 | Cons _ h t => h :: unject t
adamc@84 80 end.
adamc@84 81
adamc@84 82 Theorem inject_inverse : forall ls, unject (inject ls) = ls.
adamc@84 83 induction ls; crush.
adamc@84 84 Qed.
adamc@100 85 (* end thide *)
adamc@100 86
adam@338 87 (* EX: Implement statically checked "car"/"hd" *)
adamc@84 88
adam@283 89 (** Now let us attempt a function that is surprisingly tricky to write. In ML, the list head function raises an exception when passed an empty list. With length-indexed lists, we can rule out such invalid calls statically, and here is a first attempt at doing so. We write [???] as a placeholder for a term that we do not know how to write, not for any real Coq notation like those introduced in the previous chapter.
adamc@84 90
adamc@84 91 [[
adamc@84 92 Definition hd n (ls : ilist (S n)) : A :=
adamc@84 93 match ls with
adamc@84 94 | Nil => ???
adamc@84 95 | Cons _ h _ => h
adamc@84 96 end.
adamc@213 97
adamc@213 98 ]]
adamc@84 99
adamc@84 100 It is not clear what to write for the [Nil] case, so we are stuck before we even turn our function over to the type checker. We could try omitting the [Nil] case:
adamc@84 101
adamc@84 102 [[
adamc@84 103 Definition hd n (ls : ilist (S n)) : A :=
adamc@84 104 match ls with
adamc@84 105 | Cons _ h _ => h
adamc@84 106 end.
adam@338 107 ]]
adamc@84 108
adam@338 109 <<
adamc@84 110 Error: Non exhaustive pattern-matching: no clause found for pattern Nil
adam@338 111 >>
adamc@84 112
adam@275 113 Unlike in ML, we cannot use inexhaustive pattern matching, because there is no conception of a %\texttt{%#<tt>#Match#</tt>#%}% exception to be thrown. In fact, recent versions of Coq %\textit{%#<i>#do#</i>#%}% allow this, by implicit translation to a [match] that considers all constructors. It is educational to discover that encoding ourselves directly. We might try using an [in] clause somehow.
adamc@84 114
adamc@84 115 [[
adamc@84 116 Definition hd n (ls : ilist (S n)) : A :=
adamc@84 117 match ls in (ilist (S n)) with
adamc@84 118 | Cons _ h _ => h
adamc@84 119 end.
adamc@84 120 ]]
adamc@84 121
adam@338 122 <<
adam@338 123 Error: The reference n was not found in the current environment
adam@338 124 >>
adam@338 125
adam@338 126 In this and other cases, we feel like we want [in] clauses with type family arguments that are not variables. Unfortunately, Coq only supports variables in those positions. A completely general mechanism could only be supported with a solution to the problem of higher-order unification%~\cite{HOU}%, which is undecidable. There %\textit{%#<i>#are#</i>#%}% useful heuristics for handling non-variable indices which are gradually making their way into Coq, but we will spend some time in this and the next few chapters on effective pattern matching on dependent types using only the primitive [match] annotations.
adamc@84 127
adamc@84 128 Our final, working attempt at [hd] uses an auxiliary function and a surprising [return] annotation. *)
adamc@84 129
adamc@100 130 (* begin thide *)
adamc@84 131 Definition hd' n (ls : ilist n) :=
adamc@84 132 match ls in (ilist n) return (match n with O => unit | S _ => A end) with
adamc@84 133 | Nil => tt
adamc@84 134 | Cons _ h _ => h
adamc@84 135 end.
adamc@84 136
adam@283 137 Check hd'.
adam@283 138 (** %\vspace{-.15in}% [[
adam@283 139 hd'
adam@283 140 : forall n : nat, ilist n -> match n with
adam@283 141 | 0 => unit
adam@283 142 | S _ => A
adam@283 143 end
adam@283 144
adam@302 145 ]]
adam@302 146 *)
adam@283 147
adamc@84 148 Definition hd n (ls : ilist (S n)) : A := hd' ls.
adamc@100 149 (* end thide *)
adamc@84 150
adam@338 151 End ilist.
adam@338 152
adamc@84 153 (** We annotate our main [match] with a type that is itself a [match]. We write that the function [hd'] returns [unit] when the list is empty and returns the carried type [A] in all other cases. In the definition of [hd], we just call [hd']. Because the index of [ls] is known to be nonzero, the type checker reduces the [match] in the type of [hd'] to [A]. *)
adamc@84 154
adamc@85 155
adam@371 156 (** * The One Rule of Dependent Pattern Matching in Coq *)
adam@371 157
adam@371 158 (** The rest of this chapter will demonstrate a few other elegant applications of dependent types in Coq. Readers encountering such ideas for the first time often feel overwhelmed, concluding that there is some magic at work whereby Coq sometimes solves the halting problem for the programmer and sometimes does not, applying automated program understanding in a way far beyond what is found in conventional languages. The point of this section is to cut off that sort of thinking right now! Dependent type-checking in Coq follows just a few algorithmic rules. Chapters 10 and 12 introduce many of those rules more formally, and the main additional rule is centered on %\index{dependent pattern matching}\emph{%#<i>#dependent pattern matching#</i>#%}% of the kind we met in the previous section.
adam@371 159
adam@371 160 A dependent pattern match is a [match] expression where the type of the overall [match] is a function of the value and/or the type of the %\index{discriminee}\emph{%#<i>#discriminee#</i>#%}%, the value being matched on. In other words, the [match] type %\emph{%#<i>#depends#</i>#%}% on the discriminee.
adam@371 161
adam@371 162 When exactly will Coq accept a dependent pattern match as well-typed? Some other dependently typed languages employ fancy decision procedures to determine when programs satisfy their very expressive types. The situation in Coq is just the opposite. Only very straightforward symbolic rules are applied. Such a design choice has its drawbacks, as it forces programmers to do more work to convince the type checker of program validity. However, the great advantage of a simple type checking algorithm is that its action on %\emph{%#<i>#invalid#</i>#%}% programs is easier to understand!
adam@371 163
adam@371 164 We come now to the one rule of dependent pattern matching in Coq. A general dependent pattern match assumes this form (with unnecessary parentheses included to make the syntax easier to parse):
adam@371 165 [[
adam@371 166 match E in (T x1 ... xn) as y return U with
adam@371 167 | C z1 ... zm => B
adam@371 168 | ...
adam@371 169 end
adam@371 170 ]]
adam@371 171
adam@371 172 The discriminee is a term [E], a value in some inductive type family [T], which takes [n] arguments. An %\index{in clause}%[in] clause binds an explicit name [xi] for the [i]th argument passed to [T] in the type of [E]. An %\index{as clause}%[as] clause binds the name [y] to refer to the discriminee [E].
adam@371 173
adam@371 174 We bind these new variables [xi] and [y] so that they may be referred to in [U], a type given in the %\index{return clause}%[return] clause. The overall type of the [match] will be [U], with [E] substituted for [y], and with each [xi] substituted by the actual argument appearing in that position within [E]'s type.
adam@371 175
adam@371 176 In general, each case of a [match] may have a pattern built up in several layers from the constructors of various inductive type families. To keep this exposition simple, we will focus on patterns that are just single applications of inductive type constructors to lists of variables. Coq actually compiles the more general kind of pattern matching into this more restricted kind automatically, so understanding the typing of [match] requires understanding the typing of [match]es lowered to match one constructor at a time.
adam@371 177
adam@371 178 The last piece of the typing rule tells how to type-check a [match] case. A generic constructor application [C z1 ... zm] has some type [T x1' ... xn'], an application of the type family used in [E]'s type, probably with occurrences of the [zi] variables. From here, a simple recipe determines what type we will require for the case body [B]. The type of [B] should be [U] with the following two substitutions applied: we replace [y] (the [as] clause variable) with [C z1 ... zm], and we replace each [xi] (the [in] clause variables) with [xi']. In other words, we specialize the result type based on what we learn based on which pattern has matched the discriminee.
adam@371 179
adam@371 180 This is an exhaustive description of the ways to specify how to take advantage of which pattern has matched! No other mechanisms come into play. For instance, there is no way to specify that the types of certain free variables should be refined based on which pattern has matched. In the rest of the book, we will learn design patterns for achieving similar effects, where each technique leads to an encoding only in terms of [in], [as], and [return] clauses.
adam@371 181
adam@371 182 A few details have been omitted above. In Chapter 3, we learned that inductive type families may have both %\index{parameters}\emph{%#<i>#parameters#</i>#%}% and regular arguments. Within an [in] clause, a parameter position must have the wildcard [_] written, instead of a variable. Furthermore, recent Coq versions are adding more and more heuristics to infer dependent [match] annotations in certain conditions. The general annotation inference problem is undecidable, so there will always be serious limitations on how much work these heuristics can do. When in doubt about why a particular dependent [match] is failing to type-check, add an explicit [return] annotation! At that point, the mechanical rule sketched in this section will provide a complete account of %``%#"#what the type checker is thinking.#"#%''% Be sure to avoid the common pitfall of writing a [return] annotation that does not mention any variables bound by [in] or [as]; such a [match] will never refine typing requirements based on which pattern has matched. (One simple exception to this rule is that, when the discriminee is a variable, that same variable may be treated as if it were repeated as an [as] clause.) *)
adam@371 183
adam@371 184
adamc@85 185 (** * A Tagless Interpreter *)
adamc@85 186
adam@338 187 (** A favorite example for motivating the power of functional programming is implementation of a simple expression language interpreter. In ML and Haskell, such interpreters are often implemented using an algebraic datatype of values, where at many points it is checked that a value was built with the right constructor of the value type. With dependent types, we can implement a %\index{tagless interpreters}\textit{%#<i>#tagless#</i>#%}% interpreter that both removes this source of runtime inefficiency and gives us more confidence that our implementation is correct. *)
adamc@85 188
adamc@85 189 Inductive type : Set :=
adamc@85 190 | Nat : type
adamc@85 191 | Bool : type
adamc@85 192 | Prod : type -> type -> type.
adamc@85 193
adamc@85 194 Inductive exp : type -> Set :=
adamc@85 195 | NConst : nat -> exp Nat
adamc@85 196 | Plus : exp Nat -> exp Nat -> exp Nat
adamc@85 197 | Eq : exp Nat -> exp Nat -> exp Bool
adamc@85 198
adamc@85 199 | BConst : bool -> exp Bool
adamc@85 200 | And : exp Bool -> exp Bool -> exp Bool
adamc@85 201 | If : forall t, exp Bool -> exp t -> exp t -> exp t
adamc@85 202
adamc@85 203 | Pair : forall t1 t2, exp t1 -> exp t2 -> exp (Prod t1 t2)
adamc@85 204 | Fst : forall t1 t2, exp (Prod t1 t2) -> exp t1
adamc@85 205 | Snd : forall t1 t2, exp (Prod t1 t2) -> exp t2.
adamc@85 206
adamc@85 207 (** We have a standard algebraic datatype [type], defining a type language of naturals, booleans, and product (pair) types. Then we have the indexed inductive type [exp], where the argument to [exp] tells us the encoded type of an expression. In effect, we are defining the typing rules for expressions simultaneously with the syntax.
adamc@85 208
adamc@85 209 We can give types and expressions semantics in a new style, based critically on the chance for %\textit{%#<i>#type-level computation#</i>#%}%. *)
adamc@85 210
adamc@85 211 Fixpoint typeDenote (t : type) : Set :=
adamc@85 212 match t with
adamc@85 213 | Nat => nat
adamc@85 214 | Bool => bool
adamc@85 215 | Prod t1 t2 => typeDenote t1 * typeDenote t2
adamc@85 216 end%type.
adamc@85 217
adam@338 218 (** The [typeDenote] function compiles types of our object language into %``%#"#native#"#%''% Coq types. It is deceptively easy to implement. The only new thing we see is the [%][type] annotation, which tells Coq to parse the [match] expression using the notations associated with types. Without this annotation, the [*] would be interpreted as multiplication on naturals, rather than as the product type constructor. The token [type] is one example of an identifer bound to a %\textit{%#<i>#notation scope#</i>#%}%. In this book, we will not go into more detail on notation scopes, but the Coq manual can be consulted for more information.
adamc@85 219
adamc@85 220 We can define a function [expDenote] that is typed in terms of [typeDenote]. *)
adamc@85 221
adamc@213 222 Fixpoint expDenote t (e : exp t) : typeDenote t :=
adamc@213 223 match e with
adamc@85 224 | NConst n => n
adamc@85 225 | Plus e1 e2 => expDenote e1 + expDenote e2
adamc@85 226 | Eq e1 e2 => if eq_nat_dec (expDenote e1) (expDenote e2) then true else false
adamc@85 227
adamc@85 228 | BConst b => b
adamc@85 229 | And e1 e2 => expDenote e1 && expDenote e2
adamc@85 230 | If _ e' e1 e2 => if expDenote e' then expDenote e1 else expDenote e2
adamc@85 231
adamc@85 232 | Pair _ _ e1 e2 => (expDenote e1, expDenote e2)
adamc@85 233 | Fst _ _ e' => fst (expDenote e')
adamc@85 234 | Snd _ _ e' => snd (expDenote e')
adamc@85 235 end.
adamc@85 236
adamc@213 237 (** Despite the fancy type, the function definition is routine. In fact, it is less complicated than what we would write in ML or Haskell 98, since we do not need to worry about pushing final values in and out of an algebraic datatype. The only unusual thing is the use of an expression of the form [if E then true else false] in the [Eq] case. Remember that [eq_nat_dec] has a rich dependent type, rather than a simple boolean type. Coq's native [if] is overloaded to work on a test of any two-constructor type, so we can use [if] to build a simple boolean from the [sumbool] that [eq_nat_dec] returns.
adamc@85 238
adamc@85 239 We can implement our old favorite, a constant folding function, and prove it correct. It will be useful to write a function [pairOut] that checks if an [exp] of [Prod] type is a pair, returning its two components if so. Unsurprisingly, a first attempt leads to a type error.
adamc@85 240
adamc@85 241 [[
adamc@85 242 Definition pairOut t1 t2 (e : exp (Prod t1 t2)) : option (exp t1 * exp t2) :=
adamc@85 243 match e in (exp (Prod t1 t2)) return option (exp t1 * exp t2) with
adamc@85 244 | Pair _ _ e1 e2 => Some (e1, e2)
adamc@85 245 | _ => None
adamc@85 246 end.
adam@338 247 ]]
adamc@85 248
adam@338 249 <<
adamc@85 250 Error: The reference t2 was not found in the current environment
adam@338 251 >>
adamc@85 252
adamc@85 253 We run again into the problem of not being able to specify non-variable arguments in [in] clauses. The problem would just be hopeless without a use of an [in] clause, though, since the result type of the [match] depends on an argument to [exp]. Our solution will be to use a more general type, as we did for [hd]. First, we define a type-valued function to use in assigning a type to [pairOut]. *)
adamc@85 254
adamc@100 255 (* EX: Define a function [pairOut : forall t1 t2, exp (Prod t1 t2) -> option (exp t1 * exp t2)] *)
adamc@100 256
adamc@100 257 (* begin thide *)
adamc@85 258 Definition pairOutType (t : type) :=
adamc@85 259 match t with
adamc@85 260 | Prod t1 t2 => option (exp t1 * exp t2)
adamc@85 261 | _ => unit
adamc@85 262 end.
adamc@85 263
adamc@85 264 (** When passed a type that is a product, [pairOutType] returns our final desired type. On any other input type, [pairOutType] returns [unit], since we do not care about extracting components of non-pairs. Now we can write another helper function to provide the default behavior of [pairOut], which we will apply for inputs that are not literal pairs. *)
adamc@85 265
adamc@85 266 Definition pairOutDefault (t : type) :=
adamc@85 267 match t return (pairOutType t) with
adamc@85 268 | Prod _ _ => None
adamc@85 269 | _ => tt
adamc@85 270 end.
adamc@85 271
adamc@85 272 (** Now [pairOut] is deceptively easy to write. *)
adamc@85 273
adamc@85 274 Definition pairOut t (e : exp t) :=
adamc@85 275 match e in (exp t) return (pairOutType t) with
adamc@85 276 | Pair _ _ e1 e2 => Some (e1, e2)
adamc@85 277 | _ => pairOutDefault _
adamc@85 278 end.
adamc@100 279 (* end thide *)
adamc@85 280
adam@338 281 (** There is one important subtlety in this definition. Coq allows us to use convenient ML-style pattern matching notation, but, internally and in proofs, we see that patterns are expanded out completely, matching one level of inductive structure at a time. Thus, the default case in the [match] above expands out to one case for each constructor of [exp] besides [Pair], and the underscore in [pairOutDefault _] is resolved differently in each case. From an ML or Haskell programmer's perspective, what we have here is type inference determining which code is run (returning either [None] or [tt]), which goes beyond what is possible with type inference guiding parametric polymorphism in Hindley-Milner languages%\index{Hindley-Milner}%, but is similar to what goes on with Haskell type classes%\index{type classes}%.
adamc@85 282
adamc@213 283 With [pairOut] available, we can write [cfold] in a straightforward way. There are really no surprises beyond that Coq verifies that this code has such an expressive type, given the small annotation burden. In some places, we see that Coq's [match] annotation inference is too smart for its own good, and we have to turn that inference off by writing [return _]. *)
adamc@85 284
adamc@204 285 Fixpoint cfold t (e : exp t) : exp t :=
adamc@204 286 match e with
adamc@85 287 | NConst n => NConst n
adamc@85 288 | Plus e1 e2 =>
adamc@85 289 let e1' := cfold e1 in
adamc@85 290 let e2' := cfold e2 in
adamc@204 291 match e1', e2' return _ with
adamc@85 292 | NConst n1, NConst n2 => NConst (n1 + n2)
adamc@85 293 | _, _ => Plus e1' e2'
adamc@85 294 end
adamc@85 295 | Eq e1 e2 =>
adamc@85 296 let e1' := cfold e1 in
adamc@85 297 let e2' := cfold e2 in
adamc@204 298 match e1', e2' return _ with
adamc@85 299 | NConst n1, NConst n2 => BConst (if eq_nat_dec n1 n2 then true else false)
adamc@85 300 | _, _ => Eq e1' e2'
adamc@85 301 end
adamc@85 302
adamc@85 303 | BConst b => BConst b
adamc@85 304 | And e1 e2 =>
adamc@85 305 let e1' := cfold e1 in
adamc@85 306 let e2' := cfold e2 in
adamc@204 307 match e1', e2' return _ with
adamc@85 308 | BConst b1, BConst b2 => BConst (b1 && b2)
adamc@85 309 | _, _ => And e1' e2'
adamc@85 310 end
adamc@85 311 | If _ e e1 e2 =>
adamc@85 312 let e' := cfold e in
adamc@85 313 match e' with
adamc@85 314 | BConst true => cfold e1
adamc@85 315 | BConst false => cfold e2
adamc@85 316 | _ => If e' (cfold e1) (cfold e2)
adamc@85 317 end
adamc@85 318
adamc@85 319 | Pair _ _ e1 e2 => Pair (cfold e1) (cfold e2)
adamc@85 320 | Fst _ _ e =>
adamc@85 321 let e' := cfold e in
adamc@85 322 match pairOut e' with
adamc@85 323 | Some p => fst p
adamc@85 324 | None => Fst e'
adamc@85 325 end
adamc@85 326 | Snd _ _ e =>
adamc@85 327 let e' := cfold e in
adamc@85 328 match pairOut e' with
adamc@85 329 | Some p => snd p
adamc@85 330 | None => Snd e'
adamc@85 331 end
adamc@85 332 end.
adamc@85 333
adamc@85 334 (** The correctness theorem for [cfold] turns out to be easy to prove, once we get over one serious hurdle. *)
adamc@85 335
adamc@85 336 Theorem cfold_correct : forall t (e : exp t), expDenote e = expDenote (cfold e).
adamc@100 337 (* begin thide *)
adamc@85 338 induction e; crush.
adamc@85 339
adamc@85 340 (** The first remaining subgoal is:
adamc@85 341
adamc@85 342 [[
adamc@85 343 expDenote (cfold e1) + expDenote (cfold e2) =
adamc@85 344 expDenote
adamc@85 345 match cfold e1 with
adamc@85 346 | NConst n1 =>
adamc@85 347 match cfold e2 with
adamc@85 348 | NConst n2 => NConst (n1 + n2)
adamc@85 349 | Plus _ _ => Plus (cfold e1) (cfold e2)
adamc@85 350 | Eq _ _ => Plus (cfold e1) (cfold e2)
adamc@85 351 | BConst _ => Plus (cfold e1) (cfold e2)
adamc@85 352 | And _ _ => Plus (cfold e1) (cfold e2)
adamc@85 353 | If _ _ _ _ => Plus (cfold e1) (cfold e2)
adamc@85 354 | Pair _ _ _ _ => Plus (cfold e1) (cfold e2)
adamc@85 355 | Fst _ _ _ => Plus (cfold e1) (cfold e2)
adamc@85 356 | Snd _ _ _ => Plus (cfold e1) (cfold e2)
adamc@85 357 end
adamc@85 358 | Plus _ _ => Plus (cfold e1) (cfold e2)
adamc@85 359 | Eq _ _ => Plus (cfold e1) (cfold e2)
adamc@85 360 | BConst _ => Plus (cfold e1) (cfold e2)
adamc@85 361 | And _ _ => Plus (cfold e1) (cfold e2)
adamc@85 362 | If _ _ _ _ => Plus (cfold e1) (cfold e2)
adamc@85 363 | Pair _ _ _ _ => Plus (cfold e1) (cfold e2)
adamc@85 364 | Fst _ _ _ => Plus (cfold e1) (cfold e2)
adamc@85 365 | Snd _ _ _ => Plus (cfold e1) (cfold e2)
adamc@85 366 end
adamc@213 367
adamc@85 368 ]]
adamc@85 369
adamc@85 370 We would like to do a case analysis on [cfold e1], and we attempt that in the way that has worked so far.
adamc@85 371
adamc@85 372 [[
adamc@85 373 destruct (cfold e1).
adam@338 374 ]]
adamc@85 375
adam@338 376 <<
adamc@85 377 User error: e1 is used in hypothesis e
adam@338 378 >>
adamc@85 379
adamc@85 380 Coq gives us another cryptic error message. Like so many others, this one basically means that Coq is not able to build some proof about dependent types. It is hard to generate helpful and specific error messages for problems like this, since that would require some kind of understanding of the dependency structure of a piece of code. We will encounter many examples of case-specific tricks for recovering from errors like this one.
adamc@85 381
adam@350 382 For our current proof, we can use a tactic [dep_destruct]%\index{tactics!dep\_destruct}% defined in the book [CpdtTactics] module. General elimination/inversion of dependently typed hypotheses is undecidable, since it must be implemented with [match] expressions that have the restriction on [in] clauses that we have already discussed. The tactic [dep_destruct] makes a best effort to handle some common cases, relying upon the more primitive %\index{tactics!dependent destruction}%[dependent destruction] tactic that comes with Coq. In a future chapter, we will learn about the explicit manipulation of equality proofs that is behind [dep_destruct]'s implementation in Ltac, but for now, we treat it as a useful black box. (In Chapter 12, we will also see how [dependent destruction] forces us to make a larger philosophical commitment about our logic than we might like, and we will see some workarounds.) *)
adamc@85 383
adamc@85 384 dep_destruct (cfold e1).
adamc@85 385
adamc@85 386 (** This successfully breaks the subgoal into 5 new subgoals, one for each constructor of [exp] that could produce an [exp Nat]. Note that [dep_destruct] is successful in ruling out the other cases automatically, in effect automating some of the work that we have done manually in implementing functions like [hd] and [pairOut].
adamc@85 387
adamc@213 388 This is the only new trick we need to learn to complete the proof. We can back up and give a short, automated proof. The main inconvenience in the proof is that we cannot write a pattern that matches a [match] without including a case for every constructor of the inductive type we match over. *)
adamc@85 389
adamc@85 390 Restart.
adamc@85 391
adamc@85 392 induction e; crush;
adamc@85 393 repeat (match goal with
adamc@213 394 | [ |- context[match cfold ?E with NConst _ => _ | Plus _ _ => _
adamc@213 395 | Eq _ _ => _ | BConst _ => _ | And _ _ => _
adamc@213 396 | If _ _ _ _ => _ | Pair _ _ _ _ => _
adamc@213 397 | Fst _ _ _ => _ | Snd _ _ _ => _ end] ] =>
adamc@213 398 dep_destruct (cfold E)
adamc@213 399 | [ |- context[match pairOut (cfold ?E) with Some _ => _
adamc@213 400 | None => _ end] ] =>
adamc@213 401 dep_destruct (cfold E)
adamc@85 402 | [ |- (if ?E then _ else _) = _ ] => destruct E
adamc@85 403 end; crush).
adamc@85 404 Qed.
adamc@100 405 (* end thide *)
adamc@86 406
adamc@86 407
adam@338 408 (** * Dependently Typed Red-Black Trees *)
adamc@94 409
adam@338 410 (** Red-black trees are a favorite purely functional data structure with an interesting invariant. We can use dependent types to enforce that operations on red-black trees preserve the invariant. For simplicity, we specialize our red-black trees to represent sets of [nat]s. *)
adamc@100 411
adamc@94 412 Inductive color : Set := Red | Black.
adamc@94 413
adamc@94 414 Inductive rbtree : color -> nat -> Set :=
adamc@94 415 | Leaf : rbtree Black 0
adamc@214 416 | RedNode : forall n, rbtree Black n -> nat -> rbtree Black n -> rbtree Red n
adamc@94 417 | BlackNode : forall c1 c2 n, rbtree c1 n -> nat -> rbtree c2 n -> rbtree Black (S n).
adamc@94 418
adamc@214 419 (** A value of type [rbtree c d] is a red-black tree node whose root has color [c] and that has black depth [d]. The latter property means that there are no more than [d] black-colored nodes on any path from the root to a leaf. *)
adamc@214 420
adamc@214 421 (** At first, it can be unclear that this choice of type indices tracks any useful property. To convince ourselves, we will prove that every red-black tree is balanced. We will phrase our theorem in terms of a depth calculating function that ignores the extra information in the types. It will be useful to parameterize this function over a combining operation, so that we can re-use the same code to calculate the minimum or maximum height among all paths from root to leaf. *)
adamc@214 422
adamc@100 423 (* EX: Prove that every [rbtree] is balanced. *)
adamc@100 424
adamc@100 425 (* begin thide *)
adamc@95 426 Require Import Max Min.
adamc@95 427
adamc@95 428 Section depth.
adamc@95 429 Variable f : nat -> nat -> nat.
adamc@95 430
adamc@214 431 Fixpoint depth c n (t : rbtree c n) : nat :=
adamc@95 432 match t with
adamc@95 433 | Leaf => 0
adamc@95 434 | RedNode _ t1 _ t2 => S (f (depth t1) (depth t2))
adamc@95 435 | BlackNode _ _ _ t1 _ t2 => S (f (depth t1) (depth t2))
adamc@95 436 end.
adamc@95 437 End depth.
adamc@95 438
adam@338 439 (** Our proof of balanced-ness decomposes naturally into a lower bound and an upper bound. We prove the lower bound first. Unsurprisingly, a tree's black depth provides such a bound on the minimum path length. We use the richly typed procedure [min_dec] to do case analysis on whether [min X Y] equals [X] or [Y]. *)
adamc@214 440
adam@283 441 Check min_dec.
adam@283 442 (** %\vspace{-.15in}% [[
adam@283 443 min_dec
adam@283 444 : forall n m : nat, {min n m = n} + {min n m = m}
adam@302 445 ]]
adam@302 446 *)
adam@283 447
adamc@95 448 Theorem depth_min : forall c n (t : rbtree c n), depth min t >= n.
adamc@95 449 induction t; crush;
adamc@95 450 match goal with
adamc@95 451 | [ |- context[min ?X ?Y] ] => destruct (min_dec X Y)
adamc@95 452 end; crush.
adamc@95 453 Qed.
adamc@95 454
adamc@214 455 (** There is an analogous upper-bound theorem based on black depth. Unfortunately, a symmetric proof script does not suffice to establish it. *)
adamc@214 456
adamc@214 457 Theorem depth_max : forall c n (t : rbtree c n), depth max t <= 2 * n + 1.
adamc@214 458 induction t; crush;
adamc@214 459 match goal with
adamc@214 460 | [ |- context[max ?X ?Y] ] => destruct (max_dec X Y)
adamc@214 461 end; crush.
adamc@214 462
adamc@214 463 (** Two subgoals remain. One of them is: [[
adamc@214 464 n : nat
adamc@214 465 t1 : rbtree Black n
adamc@214 466 n0 : nat
adamc@214 467 t2 : rbtree Black n
adamc@214 468 IHt1 : depth max t1 <= n + (n + 0) + 1
adamc@214 469 IHt2 : depth max t2 <= n + (n + 0) + 1
adamc@214 470 e : max (depth max t1) (depth max t2) = depth max t1
adamc@214 471 ============================
adamc@214 472 S (depth max t1) <= n + (n + 0) + 1
adamc@214 473
adamc@214 474 ]]
adamc@214 475
adamc@214 476 We see that [IHt1] is %\textit{%#<i>#almost#</i>#%}% the fact we need, but it is not quite strong enough. We will need to strengthen our induction hypothesis to get the proof to go through. *)
adamc@214 477
adamc@214 478 Abort.
adamc@214 479
adamc@214 480 (** In particular, we prove a lemma that provides a stronger upper bound for trees with black root nodes. We got stuck above in a case about a red root node. Since red nodes have only black children, our IH strengthening will enable us to finish the proof. *)
adamc@214 481
adamc@95 482 Lemma depth_max' : forall c n (t : rbtree c n), match c with
adamc@95 483 | Red => depth max t <= 2 * n + 1
adamc@95 484 | Black => depth max t <= 2 * n
adamc@95 485 end.
adamc@95 486 induction t; crush;
adamc@95 487 match goal with
adamc@95 488 | [ |- context[max ?X ?Y] ] => destruct (max_dec X Y)
adamc@100 489 end; crush;
adamc@100 490 repeat (match goal with
adamc@214 491 | [ H : context[match ?C with Red => _ | Black => _ end] |- _ ] =>
adamc@214 492 destruct C
adamc@100 493 end; crush).
adamc@95 494 Qed.
adamc@95 495
adam@338 496 (** The original theorem follows easily from the lemma. We use the tactic %\index{tactics!generalize}%[generalize pf], which, when [pf] proves the proposition [P], changes the goal from [Q] to [P -> Q]. This transformation is useful because it makes the truth of [P] manifest syntactically, so that automation machinery can rely on [P], even if that machinery is not smart enough to establish [P] on its own. *)
adamc@214 497
adamc@95 498 Theorem depth_max : forall c n (t : rbtree c n), depth max t <= 2 * n + 1.
adamc@95 499 intros; generalize (depth_max' t); destruct c; crush.
adamc@95 500 Qed.
adamc@95 501
adamc@214 502 (** The final balance theorem establishes that the minimum and maximum path lengths of any tree are within a factor of two of each other. *)
adamc@214 503
adamc@95 504 Theorem balanced : forall c n (t : rbtree c n), 2 * depth min t + 1 >= depth max t.
adamc@95 505 intros; generalize (depth_min t); generalize (depth_max t); crush.
adamc@95 506 Qed.
adamc@100 507 (* end thide *)
adamc@95 508
adamc@214 509 (** Now we are ready to implement an example operation on our trees, insertion. Insertion can be thought of as breaking the tree invariants locally but then rebalancing. In particular, in intermediate states we find red nodes that may have red children. The type [rtree] captures the idea of such a node, continuing to track black depth as a type index. *)
adamc@95 510
adamc@94 511 Inductive rtree : nat -> Set :=
adamc@94 512 | RedNode' : forall c1 c2 n, rbtree c1 n -> nat -> rbtree c2 n -> rtree n.
adamc@94 513
adam@338 514 (** Before starting to define [insert], we define predicates capturing when a data value is in the set represented by a normal or possibly invalid tree. *)
adamc@214 515
adamc@96 516 Section present.
adamc@96 517 Variable x : nat.
adamc@96 518
adamc@214 519 Fixpoint present c n (t : rbtree c n) : Prop :=
adamc@96 520 match t with
adamc@96 521 | Leaf => False
adamc@96 522 | RedNode _ a y b => present a \/ x = y \/ present b
adamc@96 523 | BlackNode _ _ _ a y b => present a \/ x = y \/ present b
adamc@96 524 end.
adamc@96 525
adamc@96 526 Definition rpresent n (t : rtree n) : Prop :=
adamc@96 527 match t with
adamc@96 528 | RedNode' _ _ _ a y b => present a \/ x = y \/ present b
adamc@96 529 end.
adamc@96 530 End present.
adamc@96 531
adam@338 532 (** Insertion relies on two balancing operations. It will be useful to give types to these operations using a relative of the subset types from last chapter. While subset types let us pair a value with a proof about that value, here we want to pair a value with another non-proof dependently typed value. The %\index{Gallina terms!sigT}%[sigT] type fills this role. *)
adamc@214 533
adamc@100 534 Locate "{ _ : _ & _ }".
adamc@214 535 (** [[
adamc@214 536 Notation Scope
adamc@214 537 "{ x : A & P }" := sigT (fun x : A => P)
adam@302 538 ]]
adam@302 539 *)
adamc@214 540
adamc@100 541 Print sigT.
adamc@214 542 (** [[
adamc@214 543 Inductive sigT (A : Type) (P : A -> Type) : Type :=
adamc@214 544 existT : forall x : A, P x -> sigT P
adam@302 545 ]]
adam@302 546 *)
adamc@214 547
adamc@214 548 (** It will be helpful to define a concise notation for the constructor of [sigT]. *)
adamc@100 549
adamc@94 550 Notation "{< x >}" := (existT _ _ x).
adamc@94 551
adamc@214 552 (** Each balance function is used to construct a new tree whose keys include the keys of two input trees, as well as a new key. One of the two input trees may violate the red-black alternation invariant (that is, it has an [rtree] type), while the other tree is known to be valid. Crucially, the two input trees have the same black depth.
adamc@214 553
adam@338 554 A balance operation may return a tree whose root is of either color. Thus, we use a [sigT] type to package the result tree with the color of its root. Here is the definition of the first balance operation, which applies when the possibly invalid [rtree] belongs to the left of the valid [rbtree].
adam@338 555
adam@338 556 A quick word of encouragement: After writing this code, even I do not understand the precise details of how balancing works! I consulted Chris Okasaki's paper %``%#"#Red-Black Trees in a Functional Setting#"#%''~\cite{Okasaki}% and transcribed the code to use dependent types. Luckily, the details are not so important here; types alone will tell us that insertion preserves balanced-ness, and we will prove that insertion produces trees containing the right keys.*)
adamc@214 557
adamc@94 558 Definition balance1 n (a : rtree n) (data : nat) c2 :=
adamc@214 559 match a in rtree n return rbtree c2 n
adamc@214 560 -> { c : color & rbtree c (S n) } with
adam@380 561 | RedNode' _ c0 _ t1 y t2 =>
adam@380 562 match t1 in rbtree c n return rbtree c0 n -> rbtree c2 n
adamc@214 563 -> { c : color & rbtree c (S n) } with
adamc@214 564 | RedNode _ a x b => fun c d =>
adamc@214 565 {<RedNode (BlackNode a x b) y (BlackNode c data d)>}
adamc@94 566 | t1' => fun t2 =>
adam@380 567 match t2 in rbtree c n return rbtree Black n -> rbtree c2 n
adamc@214 568 -> { c : color & rbtree c (S n) } with
adamc@214 569 | RedNode _ b x c => fun a d =>
adamc@214 570 {<RedNode (BlackNode a y b) x (BlackNode c data d)>}
adamc@95 571 | b => fun a t => {<BlackNode (RedNode a y b) data t>}
adamc@94 572 end t1'
adamc@94 573 end t2
adamc@94 574 end.
adamc@94 575
adam@338 576 (** We apply a trick that I call the %\index{convoy pattern}\textit{%#<i>#convoy pattern#</i>#%}%. Recall that [match] annotations only make it possible to describe a dependence of a [match] %\textit{%#<i>#result type#</i>#%}% on the discriminee. There is no automatic refinement of the types of free variables. However, it is possible to effect such a refinement by finding a way to encode free variable type dependencies in the [match] result type, so that a [return] clause can express the connection.
adamc@214 577
adam@292 578 In particular, we can extend the [match] to return %\textit{%#<i>#functions over the free variables whose types we want to refine#</i>#%}%. In the case of [balance1], we only find ourselves wanting to refine the type of one tree variable at a time. We match on one subtree of a node, and we want the type of the other subtree to be refined based on what we learn. We indicate this with a [return] clause starting like [rbtree _ n -> ...], where [n] is bound in an [in] pattern. Such a [match] expression is applied immediately to the %``%#"#old version#"#%''% of the variable to be refined, and the type checker is happy.
adamc@214 579
adam@338 580 Here is the symmetric function [balance2], for cases where the possibly invalid tree appears on the right rather than on the left. *)
adamc@214 581
adamc@94 582 Definition balance2 n (a : rtree n) (data : nat) c2 :=
adamc@94 583 match a in rtree n return rbtree c2 n -> { c : color & rbtree c (S n) } with
adam@380 584 | RedNode' _ c0 _ t1 z t2 =>
adam@380 585 match t1 in rbtree c n return rbtree c0 n -> rbtree c2 n
adamc@214 586 -> { c : color & rbtree c (S n) } with
adamc@214 587 | RedNode _ b y c => fun d a =>
adamc@214 588 {<RedNode (BlackNode a data b) y (BlackNode c z d)>}
adamc@94 589 | t1' => fun t2 =>
adam@380 590 match t2 in rbtree c n return rbtree Black n -> rbtree c2 n
adamc@214 591 -> { c : color & rbtree c (S n) } with
adamc@214 592 | RedNode _ c z' d => fun b a =>
adamc@214 593 {<RedNode (BlackNode a data b) z (BlackNode c z' d)>}
adamc@95 594 | b => fun a t => {<BlackNode t data (RedNode a z b)>}
adamc@94 595 end t1'
adamc@94 596 end t2
adamc@94 597 end.
adamc@94 598
adamc@214 599 (** Now we are almost ready to get down to the business of writing an [insert] function. First, we enter a section that declares a variable [x], for the key we want to insert. *)
adamc@214 600
adamc@94 601 Section insert.
adamc@94 602 Variable x : nat.
adamc@94 603
adamc@214 604 (** Most of the work of insertion is done by a helper function [ins], whose return types are expressed using a type-level function [insResult]. *)
adamc@214 605
adamc@94 606 Definition insResult c n :=
adamc@94 607 match c with
adamc@94 608 | Red => rtree n
adamc@94 609 | Black => { c' : color & rbtree c' n }
adamc@94 610 end.
adamc@94 611
adam@338 612 (** That is, inserting into a tree with root color [c] and black depth [n], the variety of tree we get out depends on [c]. If we started with a red root, then we get back a possibly invalid tree of depth [n]. If we started with a black root, we get back a valid tree of depth [n] with a root node of an arbitrary color.
adamc@214 613
adamc@214 614 Here is the definition of [ins]. Again, we do not want to dwell on the functional details. *)
adamc@214 615
adamc@214 616 Fixpoint ins c n (t : rbtree c n) : insResult c n :=
adamc@214 617 match t with
adamc@94 618 | Leaf => {< RedNode Leaf x Leaf >}
adamc@94 619 | RedNode _ a y b =>
adamc@94 620 if le_lt_dec x y
adamc@94 621 then RedNode' (projT2 (ins a)) y b
adamc@94 622 else RedNode' a y (projT2 (ins b))
adamc@94 623 | BlackNode c1 c2 _ a y b =>
adamc@94 624 if le_lt_dec x y
adamc@94 625 then
adamc@94 626 match c1 return insResult c1 _ -> _ with
adamc@94 627 | Red => fun ins_a => balance1 ins_a y b
adamc@94 628 | _ => fun ins_a => {< BlackNode (projT2 ins_a) y b >}
adamc@94 629 end (ins a)
adamc@94 630 else
adamc@94 631 match c2 return insResult c2 _ -> _ with
adamc@94 632 | Red => fun ins_b => balance2 ins_b y a
adamc@94 633 | _ => fun ins_b => {< BlackNode a y (projT2 ins_b) >}
adamc@94 634 end (ins b)
adamc@94 635 end.
adamc@94 636
adam@338 637 (** The one new trick is a variation of the convoy pattern. In each of the last two pattern matches, we want to take advantage of the typing connection between the trees [a] and [b]. We might naively apply the convoy pattern directly on [a] in the first [match] and on [b] in the second. This satisfies the type checker per se, but it does not satisfy the termination checker. Inside each [match], we would be calling [ins] recursively on a locally bound variable. The termination checker is not smart enough to trace the dataflow into that variable, so the checker does not know that this recursive argument is smaller than the original argument. We make this fact clearer by applying the convoy pattern on %\textit{%#<i>#the result of a recursive call#</i>#%}%, rather than just on that call's argument.
adamc@214 638
adamc@214 639 Finally, we are in the home stretch of our effort to define [insert]. We just need a few more definitions of non-recursive functions. First, we need to give the final characterization of [insert]'s return type. Inserting into a red-rooted tree gives a black-rooted tree where black depth has increased, and inserting into a black-rooted tree gives a tree where black depth has stayed the same and where the root is an arbitrary color. *)
adamc@214 640
adamc@94 641 Definition insertResult c n :=
adamc@94 642 match c with
adamc@94 643 | Red => rbtree Black (S n)
adamc@94 644 | Black => { c' : color & rbtree c' n }
adamc@94 645 end.
adamc@94 646
adamc@214 647 (** A simple clean-up procedure translates [insResult]s into [insertResult]s. *)
adamc@214 648
adamc@97 649 Definition makeRbtree c n : insResult c n -> insertResult c n :=
adamc@214 650 match c with
adamc@94 651 | Red => fun r =>
adamc@214 652 match r with
adamc@94 653 | RedNode' _ _ _ a x b => BlackNode a x b
adamc@94 654 end
adamc@94 655 | Black => fun r => r
adamc@94 656 end.
adamc@94 657
adamc@214 658 (** We modify Coq's default choice of implicit arguments for [makeRbtree], so that we do not need to specify the [c] and [n] arguments explicitly in later calls. *)
adamc@214 659
adamc@97 660 Implicit Arguments makeRbtree [c n].
adamc@94 661
adamc@214 662 (** Finally, we define [insert] as a simple composition of [ins] and [makeRbtree]. *)
adamc@214 663
adamc@94 664 Definition insert c n (t : rbtree c n) : insertResult c n :=
adamc@97 665 makeRbtree (ins t).
adamc@94 666
adamc@214 667 (** As we noted earlier, the type of [insert] guarantees that it outputs balanced trees whose depths have not increased too much. We also want to know that [insert] operates correctly on trees interpreted as finite sets, so we finish this section with a proof of that fact. *)
adamc@214 668
adamc@95 669 Section present.
adamc@95 670 Variable z : nat.
adamc@95 671
adamc@214 672 (** The variable [z] stands for an arbitrary key. We will reason about [z]'s presence in particular trees. As usual, outside the section the theorems we prove will quantify over all possible keys, giving us the facts we wanted.
adamc@214 673
adam@367 674 We start by proving the correctness of the balance operations. It is useful to define a custom tactic [present_balance] that encapsulates the reasoning common to the two proofs. We use the keyword %\index{Vernacular commands!Ltac}%[Ltac] to assign a name to a proof script. This particular script just iterates between [crush] and identification of a tree that is being pattern-matched on and should be destructed. *)
adamc@214 675
adamc@98 676 Ltac present_balance :=
adamc@98 677 crush;
adamc@98 678 repeat (match goal with
adam@338 679 | [ _ : context[match ?T with
adamc@98 680 | Leaf => _
adamc@98 681 | RedNode _ _ _ _ => _
adamc@98 682 | BlackNode _ _ _ _ _ _ => _
adamc@98 683 end] |- _ ] => dep_destruct T
adamc@98 684 | [ |- context[match ?T with
adamc@98 685 | Leaf => _
adamc@98 686 | RedNode _ _ _ _ => _
adamc@98 687 | BlackNode _ _ _ _ _ _ => _
adamc@98 688 end] ] => dep_destruct T
adamc@98 689 end; crush).
adamc@98 690
adamc@214 691 (** The balance correctness theorems are simple first-order logic equivalences, where we use the function [projT2] to project the payload of a [sigT] value. *)
adamc@214 692
adam@294 693 Lemma present_balance1 : forall n (a : rtree n) (y : nat) c2 (b : rbtree c2 n),
adamc@95 694 present z (projT2 (balance1 a y b))
adamc@95 695 <-> rpresent z a \/ z = y \/ present z b.
adamc@98 696 destruct a; present_balance.
adamc@95 697 Qed.
adamc@95 698
adamc@213 699 Lemma present_balance2 : forall n (a : rtree n) (y : nat) c2 (b : rbtree c2 n),
adamc@95 700 present z (projT2 (balance2 a y b))
adamc@95 701 <-> rpresent z a \/ z = y \/ present z b.
adamc@98 702 destruct a; present_balance.
adamc@95 703 Qed.
adamc@95 704
adamc@214 705 (** To state the theorem for [ins], it is useful to define a new type-level function, since [ins] returns different result types based on the type indices passed to it. Recall that [x] is the section variable standing for the key we are inserting. *)
adamc@214 706
adamc@95 707 Definition present_insResult c n :=
adamc@95 708 match c return (rbtree c n -> insResult c n -> Prop) with
adamc@95 709 | Red => fun t r => rpresent z r <-> z = x \/ present z t
adamc@95 710 | Black => fun t r => present z (projT2 r) <-> z = x \/ present z t
adamc@95 711 end.
adamc@95 712
adamc@214 713 (** Now the statement and proof of the [ins] correctness theorem are straightforward, if verbose. We proceed by induction on the structure of a tree, followed by finding case analysis opportunities on expressions we see being analyzed in [if] or [match] expressions. After that, we pattern-match to find opportunities to use the theorems we proved about balancing. Finally, we identify two variables that are asserted by some hypothesis to be equal, and we use that hypothesis to replace one variable with the other everywhere. *)
adamc@214 714
adamc@214 715 (** printing * $*$ *)
adamc@214 716
adamc@95 717 Theorem present_ins : forall c n (t : rbtree c n),
adamc@95 718 present_insResult t (ins t).
adamc@95 719 induction t; crush;
adamc@95 720 repeat (match goal with
adam@338 721 | [ _ : context[if ?E then _ else _] |- _ ] => destruct E
adamc@95 722 | [ |- context[if ?E then _ else _] ] => destruct E
adam@338 723 | [ _ : context[match ?C with Red => _ | Black => _ end]
adamc@214 724 |- _ ] => destruct C
adamc@95 725 end; crush);
adamc@95 726 try match goal with
adam@338 727 | [ _ : context[balance1 ?A ?B ?C] |- _ ] =>
adamc@95 728 generalize (present_balance1 A B C)
adamc@95 729 end;
adamc@95 730 try match goal with
adam@338 731 | [ _ : context[balance2 ?A ?B ?C] |- _ ] =>
adamc@95 732 generalize (present_balance2 A B C)
adamc@95 733 end;
adamc@95 734 try match goal with
adamc@95 735 | [ |- context[balance1 ?A ?B ?C] ] =>
adamc@95 736 generalize (present_balance1 A B C)
adamc@95 737 end;
adamc@95 738 try match goal with
adamc@95 739 | [ |- context[balance2 ?A ?B ?C] ] =>
adamc@95 740 generalize (present_balance2 A B C)
adamc@95 741 end;
adamc@214 742 crush;
adamc@95 743 match goal with
adamc@95 744 | [ z : nat, x : nat |- _ ] =>
adamc@95 745 match goal with
adamc@95 746 | [ H : z = x |- _ ] => rewrite H in *; clear H
adamc@95 747 end
adamc@95 748 end;
adamc@95 749 tauto.
adamc@95 750 Qed.
adamc@95 751
adamc@214 752 (** printing * $\times$ *)
adamc@214 753
adamc@214 754 (** The hard work is done. The most readable way to state correctness of [insert] involves splitting the property into two color-specific theorems. We write a tactic to encapsulate the reasoning steps that work to establish both facts. *)
adamc@214 755
adamc@213 756 Ltac present_insert :=
adamc@213 757 unfold insert; intros n t; inversion t;
adamc@97 758 generalize (present_ins t); simpl;
adamc@97 759 dep_destruct (ins t); tauto.
adamc@97 760
adamc@95 761 Theorem present_insert_Red : forall n (t : rbtree Red n),
adamc@95 762 present z (insert t)
adamc@95 763 <-> (z = x \/ present z t).
adamc@213 764 present_insert.
adamc@95 765 Qed.
adamc@95 766
adamc@95 767 Theorem present_insert_Black : forall n (t : rbtree Black n),
adamc@95 768 present z (projT2 (insert t))
adamc@95 769 <-> (z = x \/ present z t).
adamc@213 770 present_insert.
adamc@95 771 Qed.
adamc@95 772 End present.
adamc@94 773 End insert.
adamc@94 774
adam@338 775 (** We can generate executable OCaml code with the command %\index{Vernacular commands!Recursive Extraction}%[Recursive Extraction insert], which also automatically outputs the OCaml versions of all of [insert]'s dependencies. In our previous extractions, we wound up with clean OCaml code. Here, we find uses of %\index{Obj.magic}\texttt{%#<tt>#Obj.magic#</tt>#%}%, OCaml's unsafe cast operator for tweaking the apparent type of an expression in an arbitrary way. Casts appear for this example because the return type of [insert] depends on the %\textit{%#<i>#value#</i>#%}% of the function's argument, a pattern which OCaml cannot handle. Since Coq's type system is much more expressive than OCaml's, such casts are unavoidable in general. Since the OCaml type-checker is no longer checking full safety of programs, we must rely on Coq's extractor to use casts only in provably safe ways. *)
adam@338 776
adam@338 777 (* begin hide *)
adam@338 778 Recursive Extraction insert.
adam@338 779 (* end hide *)
adam@283 780
adamc@94 781
adamc@86 782 (** * A Certified Regular Expression Matcher *)
adamc@86 783
adamc@93 784 (** Another interesting example is regular expressions with dependent types that express which predicates over strings particular regexps implement. We can then assign a dependent type to a regular expression matching function, guaranteeing that it always decides the string property that we expect it to decide.
adamc@93 785
adam@338 786 Before defining the syntax of expressions, it is helpful to define an inductive type capturing the meaning of the Kleene star. That is, a string [s] matches regular expression [star e] if and only if [s] can be decomposed into a sequence of substrings that all match [e]. We use Coq's string support, which comes through a combination of the [Strings] library and some parsing notations built into Coq. Operators like [++] and functions like [length] that we know from lists are defined again for strings. Notation scopes help us control which versions we want to use in particular contexts.%\index{Vernacular commands!Open Scope}% *)
adamc@93 787
adamc@86 788 Require Import Ascii String.
adamc@86 789 Open Scope string_scope.
adamc@86 790
adamc@91 791 Section star.
adamc@91 792 Variable P : string -> Prop.
adamc@91 793
adamc@91 794 Inductive star : string -> Prop :=
adamc@91 795 | Empty : star ""
adamc@91 796 | Iter : forall s1 s2,
adamc@91 797 P s1
adamc@91 798 -> star s2
adamc@91 799 -> star (s1 ++ s2).
adamc@91 800 End star.
adamc@91 801
adam@283 802 (** Now we can make our first attempt at defining a [regexp] type that is indexed by predicates on strings. Here is a reasonable-looking definition that is restricted to constant characters and concatenation. We use the constructor [String], which is the analogue of list cons for the type [string], where [""] is like list nil.
adamc@93 803 [[
adamc@93 804 Inductive regexp : (string -> Prop) -> Set :=
adamc@93 805 | Char : forall ch : ascii,
adamc@93 806 regexp (fun s => s = String ch "")
adamc@93 807 | Concat : forall (P1 P2 : string -> Prop) (r1 : regexp P1) (r2 : regexp P2),
adamc@93 808 regexp (fun s => exists s1, exists s2, s = s1 ++ s2 /\ P1 s1 /\ P2 s2).
adamc@93 809 ]]
adamc@93 810
adam@338 811 <<
adam@338 812 User error: Large non-propositional inductive types must be in Type
adam@338 813 >>
adam@338 814
adam@338 815 What is a %\index{large inductive types}%large inductive type? In Coq, it is an inductive type that has a constructor which quantifies over some type of type [Type]. We have not worked with [Type] very much to this point. Every term of CIC has a type, including [Set] and [Prop], which are assigned type [Type]. The type [string -> Prop] from the failed definition also has type [Type].
adamc@93 816
adamc@93 817 It turns out that allowing large inductive types in [Set] leads to contradictions when combined with certain kinds of classical logic reasoning. Thus, by default, such types are ruled out. There is a simple fix for our [regexp] definition, which is to place our new type in [Type]. While fixing the problem, we also expand the list of constructors to cover the remaining regular expression operators. *)
adamc@93 818
adamc@89 819 Inductive regexp : (string -> Prop) -> Type :=
adamc@86 820 | Char : forall ch : ascii,
adamc@86 821 regexp (fun s => s = String ch "")
adamc@86 822 | Concat : forall P1 P2 (r1 : regexp P1) (r2 : regexp P2),
adamc@87 823 regexp (fun s => exists s1, exists s2, s = s1 ++ s2 /\ P1 s1 /\ P2 s2)
adamc@87 824 | Or : forall P1 P2 (r1 : regexp P1) (r2 : regexp P2),
adamc@91 825 regexp (fun s => P1 s \/ P2 s)
adamc@91 826 | Star : forall P (r : regexp P),
adamc@91 827 regexp (star P).
adamc@86 828
adam@296 829 (** Many theorems about strings are useful for implementing a certified regexp matcher, and few of them are in the [Strings] library. The book source includes statements, proofs, and hint commands for a handful of such omitted theorems. Since they are orthogonal to our use of dependent types, we hide them in the rendered versions of this book. *)
adamc@93 830
adamc@93 831 (* begin hide *)
adamc@86 832 Open Scope specif_scope.
adamc@86 833
adamc@86 834 Lemma length_emp : length "" <= 0.
adamc@86 835 crush.
adamc@86 836 Qed.
adamc@86 837
adamc@86 838 Lemma append_emp : forall s, s = "" ++ s.
adamc@86 839 crush.
adamc@86 840 Qed.
adamc@86 841
adamc@86 842 Ltac substring :=
adamc@86 843 crush;
adamc@86 844 repeat match goal with
adamc@86 845 | [ |- context[match ?N with O => _ | S _ => _ end] ] => destruct N; crush
adamc@86 846 end.
adamc@86 847
adamc@86 848 Lemma substring_le : forall s n m,
adamc@86 849 length (substring n m s) <= m.
adamc@86 850 induction s; substring.
adamc@86 851 Qed.
adamc@86 852
adamc@86 853 Lemma substring_all : forall s,
adamc@86 854 substring 0 (length s) s = s.
adamc@86 855 induction s; substring.
adamc@86 856 Qed.
adamc@86 857
adamc@86 858 Lemma substring_none : forall s n,
adamc@93 859 substring n 0 s = "".
adamc@86 860 induction s; substring.
adamc@86 861 Qed.
adamc@86 862
adam@375 863 Hint Rewrite substring_all substring_none.
adamc@86 864
adamc@86 865 Lemma substring_split : forall s m,
adamc@86 866 substring 0 m s ++ substring m (length s - m) s = s.
adamc@86 867 induction s; substring.
adamc@86 868 Qed.
adamc@86 869
adamc@86 870 Lemma length_app1 : forall s1 s2,
adamc@86 871 length s1 <= length (s1 ++ s2).
adamc@86 872 induction s1; crush.
adamc@86 873 Qed.
adamc@86 874
adamc@86 875 Hint Resolve length_emp append_emp substring_le substring_split length_app1.
adamc@86 876
adamc@86 877 Lemma substring_app_fst : forall s2 s1 n,
adamc@86 878 length s1 = n
adamc@86 879 -> substring 0 n (s1 ++ s2) = s1.
adamc@86 880 induction s1; crush.
adamc@86 881 Qed.
adamc@86 882
adamc@86 883 Lemma substring_app_snd : forall s2 s1 n,
adamc@86 884 length s1 = n
adamc@86 885 -> substring n (length (s1 ++ s2) - n) (s1 ++ s2) = s2.
adam@375 886 Hint Rewrite <- minus_n_O.
adamc@86 887
adamc@86 888 induction s1; crush.
adamc@86 889 Qed.
adamc@86 890
adam@375 891 Hint Rewrite substring_app_fst substring_app_snd using solve [trivial].
adamc@93 892 (* end hide *)
adamc@93 893
adamc@93 894 (** A few auxiliary functions help us in our final matcher definition. The function [split] will be used to implement the regexp concatenation case. *)
adamc@86 895
adamc@86 896 Section split.
adamc@86 897 Variables P1 P2 : string -> Prop.
adamc@214 898 Variable P1_dec : forall s, {P1 s} + {~ P1 s}.
adamc@214 899 Variable P2_dec : forall s, {P2 s} + {~ P2 s}.
adamc@93 900 (** We require a choice of two arbitrary string predicates and functions for deciding them. *)
adamc@86 901
adamc@86 902 Variable s : string.
adamc@93 903 (** Our computation will take place relative to a single fixed string, so it is easiest to make it a [Variable], rather than an explicit argument to our functions. *)
adamc@93 904
adam@338 905 (** The function [split'] is the workhorse behind [split]. It searches through the possible ways of splitting [s] into two pieces, checking the two predicates against each such pair. The execution of [split'] progresses right-to-left, from splitting all of [s] into the first piece to splitting all of [s] into the second piece. It takes an extra argument, [n], which specifies how far along we are in this search process. *)
adamc@86 906
adam@297 907 Definition split' : forall n : nat, n <= length s
adamc@86 908 -> {exists s1, exists s2, length s1 <= n /\ s1 ++ s2 = s /\ P1 s1 /\ P2 s2}
adamc@214 909 + {forall s1 s2, length s1 <= n -> s1 ++ s2 = s -> ~ P1 s1 \/ ~ P2 s2}.
adamc@86 910 refine (fix F (n : nat) : n <= length s
adamc@86 911 -> {exists s1, exists s2, length s1 <= n /\ s1 ++ s2 = s /\ P1 s1 /\ P2 s2}
adamc@214 912 + {forall s1 s2, length s1 <= n -> s1 ++ s2 = s -> ~ P1 s1 \/ ~ P2 s2} :=
adamc@214 913 match n with
adamc@86 914 | O => fun _ => Reduce (P1_dec "" && P2_dec s)
adamc@93 915 | S n' => fun _ => (P1_dec (substring 0 (S n') s)
adamc@93 916 && P2_dec (substring (S n') (length s - S n') s))
adamc@86 917 || F n' _
adamc@86 918 end); clear F; crush; eauto 7;
adamc@86 919 match goal with
adamc@86 920 | [ _ : length ?S <= 0 |- _ ] => destruct S
adam@338 921 | [ _ : length ?S' <= S ?N |- _ ] => destruct (eq_nat_dec (length S') (S N))
adamc@86 922 end; crush.
adamc@86 923 Defined.
adamc@86 924
adam@338 925 (** There is one subtle point in the [split'] code that is worth mentioning. The main body of the function is a [match] on [n]. In the case where [n] is known to be [S n'], we write [S n'] in several places where we might be tempted to write [n]. However, without further work to craft proper [match] annotations, the type-checker does not use the equality between [n] and [S n']. Thus, it is common to see patterns repeated in [match] case bodies in dependently typed Coq code. We can at least use a [let] expression to avoid copying the pattern more than once, replacing the first case body with:
adamc@93 926 [[
adamc@93 927 | S n' => fun _ => let n := S n' in
adamc@93 928 (P1_dec (substring 0 n s)
adamc@93 929 && P2_dec (substring n (length s - n) s))
adamc@93 930 || F n' _
adamc@214 931
adamc@93 932 ]]
adamc@93 933
adam@338 934 The [split] function itself is trivial to implement in terms of [split']. We just ask [split'] to begin its search with [n = length s]. *)
adamc@93 935
adamc@86 936 Definition split : {exists s1, exists s2, s = s1 ++ s2 /\ P1 s1 /\ P2 s2}
adamc@214 937 + {forall s1 s2, s = s1 ++ s2 -> ~ P1 s1 \/ ~ P2 s2}.
adamc@86 938 refine (Reduce (split' (n := length s) _)); crush; eauto.
adamc@86 939 Defined.
adamc@86 940 End split.
adamc@86 941
adamc@86 942 Implicit Arguments split [P1 P2].
adamc@86 943
adamc@93 944 (* begin hide *)
adamc@91 945 Lemma app_empty_end : forall s, s ++ "" = s.
adamc@91 946 induction s; crush.
adamc@91 947 Qed.
adamc@91 948
adam@375 949 Hint Rewrite app_empty_end.
adamc@91 950
adamc@91 951 Lemma substring_self : forall s n,
adamc@91 952 n <= 0
adamc@91 953 -> substring n (length s - n) s = s.
adamc@91 954 induction s; substring.
adamc@91 955 Qed.
adamc@91 956
adamc@91 957 Lemma substring_empty : forall s n m,
adamc@91 958 m <= 0
adamc@91 959 -> substring n m s = "".
adamc@91 960 induction s; substring.
adamc@91 961 Qed.
adamc@91 962
adam@375 963 Hint Rewrite substring_self substring_empty using omega.
adamc@91 964
adamc@91 965 Lemma substring_split' : forall s n m,
adamc@91 966 substring n m s ++ substring (n + m) (length s - (n + m)) s
adamc@91 967 = substring n (length s - n) s.
adam@375 968 Hint Rewrite substring_split.
adamc@91 969
adamc@91 970 induction s; substring.
adamc@91 971 Qed.
adamc@91 972
adamc@91 973 Lemma substring_stack : forall s n2 m1 m2,
adamc@91 974 m1 <= m2
adamc@91 975 -> substring 0 m1 (substring n2 m2 s)
adamc@91 976 = substring n2 m1 s.
adamc@91 977 induction s; substring.
adamc@91 978 Qed.
adamc@91 979
adamc@91 980 Ltac substring' :=
adamc@91 981 crush;
adamc@91 982 repeat match goal with
adamc@91 983 | [ |- context[match ?N with O => _ | S _ => _ end] ] => case_eq N; crush
adamc@91 984 end.
adamc@91 985
adamc@91 986 Lemma substring_stack' : forall s n1 n2 m1 m2,
adamc@91 987 n1 + m1 <= m2
adamc@91 988 -> substring n1 m1 (substring n2 m2 s)
adamc@91 989 = substring (n1 + n2) m1 s.
adamc@91 990 induction s; substring';
adamc@91 991 match goal with
adamc@91 992 | [ |- substring ?N1 _ _ = substring ?N2 _ _ ] =>
adamc@91 993 replace N1 with N2; crush
adamc@91 994 end.
adamc@91 995 Qed.
adamc@91 996
adamc@91 997 Lemma substring_suffix : forall s n,
adamc@91 998 n <= length s
adamc@91 999 -> length (substring n (length s - n) s) = length s - n.
adamc@91 1000 induction s; substring.
adamc@91 1001 Qed.
adamc@91 1002
adamc@91 1003 Lemma substring_suffix_emp' : forall s n m,
adamc@91 1004 substring n (S m) s = ""
adamc@91 1005 -> n >= length s.
adamc@91 1006 induction s; crush;
adamc@91 1007 match goal with
adamc@91 1008 | [ |- ?N >= _ ] => destruct N; crush
adamc@91 1009 end;
adamc@91 1010 match goal with
adamc@91 1011 [ |- S ?N >= S ?E ] => assert (N >= E); [ eauto | omega ]
adamc@91 1012 end.
adamc@91 1013 Qed.
adamc@91 1014
adamc@91 1015 Lemma substring_suffix_emp : forall s n m,
adamc@92 1016 substring n m s = ""
adamc@92 1017 -> m > 0
adamc@91 1018 -> n >= length s.
adam@335 1019 destruct m as [ | m]; [crush | intros; apply substring_suffix_emp' with m; assumption].
adamc@91 1020 Qed.
adamc@91 1021
adamc@91 1022 Hint Rewrite substring_stack substring_stack' substring_suffix
adam@375 1023 using omega.
adamc@91 1024
adamc@91 1025 Lemma minus_minus : forall n m1 m2,
adamc@91 1026 m1 + m2 <= n
adamc@91 1027 -> n - m1 - m2 = n - (m1 + m2).
adamc@91 1028 intros; omega.
adamc@91 1029 Qed.
adamc@91 1030
adamc@91 1031 Lemma plus_n_Sm' : forall n m : nat, S (n + m) = m + S n.
adamc@91 1032 intros; omega.
adamc@91 1033 Qed.
adamc@91 1034
adam@375 1035 Hint Rewrite minus_minus using omega.
adamc@93 1036 (* end hide *)
adamc@93 1037
adamc@93 1038 (** One more helper function will come in handy: [dec_star], for implementing another linear search through ways of splitting a string, this time for implementing the Kleene star. *)
adamc@91 1039
adamc@91 1040 Section dec_star.
adamc@91 1041 Variable P : string -> Prop.
adamc@214 1042 Variable P_dec : forall s, {P s} + {~ P s}.
adamc@91 1043
adam@338 1044 (** Some new lemmas and hints about the [star] type family are useful. We omit them here; they are included in the book source at this point. *)
adamc@93 1045
adamc@93 1046 (* begin hide *)
adamc@91 1047 Hint Constructors star.
adamc@91 1048
adamc@91 1049 Lemma star_empty : forall s,
adamc@91 1050 length s = 0
adamc@91 1051 -> star P s.
adamc@91 1052 destruct s; crush.
adamc@91 1053 Qed.
adamc@91 1054
adamc@91 1055 Lemma star_singleton : forall s, P s -> star P s.
adamc@91 1056 intros; rewrite <- (app_empty_end s); auto.
adamc@91 1057 Qed.
adamc@91 1058
adamc@91 1059 Lemma star_app : forall s n m,
adamc@91 1060 P (substring n m s)
adamc@91 1061 -> star P (substring (n + m) (length s - (n + m)) s)
adamc@91 1062 -> star P (substring n (length s - n) s).
adamc@91 1063 induction n; substring;
adamc@91 1064 match goal with
adamc@91 1065 | [ H : P (substring ?N ?M ?S) |- _ ] =>
adamc@91 1066 solve [ rewrite <- (substring_split S M); auto
adamc@91 1067 | rewrite <- (substring_split' S N M); auto ]
adamc@91 1068 end.
adamc@91 1069 Qed.
adamc@91 1070
adamc@91 1071 Hint Resolve star_empty star_singleton star_app.
adamc@91 1072
adamc@91 1073 Variable s : string.
adamc@91 1074
adamc@91 1075 Lemma star_inv : forall s,
adamc@91 1076 star P s
adamc@91 1077 -> s = ""
adamc@91 1078 \/ exists i, i < length s
adamc@91 1079 /\ P (substring 0 (S i) s)
adamc@91 1080 /\ star P (substring (S i) (length s - S i) s).
adamc@91 1081 Hint Extern 1 (exists i : nat, _) =>
adamc@91 1082 match goal with
adamc@91 1083 | [ H : P (String _ ?S) |- _ ] => exists (length S); crush
adamc@91 1084 end.
adamc@91 1085
adamc@91 1086 induction 1; [
adamc@91 1087 crush
adamc@91 1088 | match goal with
adamc@91 1089 | [ _ : P ?S |- _ ] => destruct S; crush
adamc@91 1090 end
adamc@91 1091 ].
adamc@91 1092 Qed.
adamc@91 1093
adamc@91 1094 Lemma star_substring_inv : forall n,
adamc@91 1095 n <= length s
adamc@91 1096 -> star P (substring n (length s - n) s)
adamc@91 1097 -> substring n (length s - n) s = ""
adamc@91 1098 \/ exists l, l < length s - n
adamc@91 1099 /\ P (substring n (S l) s)
adamc@91 1100 /\ star P (substring (n + S l) (length s - (n + S l)) s).
adam@375 1101 Hint Rewrite plus_n_Sm'.
adamc@91 1102
adamc@91 1103 intros;
adamc@91 1104 match goal with
adamc@91 1105 | [ H : star _ _ |- _ ] => generalize (star_inv H); do 3 crush; eauto
adamc@91 1106 end.
adamc@91 1107 Qed.
adamc@93 1108 (* end hide *)
adamc@93 1109
adamc@93 1110 (** The function [dec_star''] implements a single iteration of the star. That is, it tries to find a string prefix matching [P], and it calls a parameter function on the remainder of the string. *)
adamc@91 1111
adamc@91 1112 Section dec_star''.
adamc@91 1113 Variable n : nat.
adamc@93 1114 (** [n] is the length of the prefix of [s] that we have already processed. *)
adamc@91 1115
adamc@91 1116 Variable P' : string -> Prop.
adamc@91 1117 Variable P'_dec : forall n' : nat, n' > n
adamc@91 1118 -> {P' (substring n' (length s - n') s)}
adamc@214 1119 + {~ P' (substring n' (length s - n') s)}.
adamc@93 1120 (** When we use [dec_star''], we will instantiate [P'_dec] with a function for continuing the search for more instances of [P] in [s]. *)
adamc@93 1121
adamc@93 1122 (** Now we come to [dec_star''] itself. It takes as an input a natural [l] that records how much of the string has been searched so far, as we did for [split']. The return type expresses that [dec_star''] is looking for an index into [s] that splits [s] into a nonempty prefix and a suffix, such that the prefix satisfies [P] and the suffix satisfies [P']. *)
adamc@91 1123
adam@297 1124 Definition dec_star'' : forall l : nat,
adam@297 1125 {exists l', S l' <= l
adamc@91 1126 /\ P (substring n (S l') s) /\ P' (substring (n + S l') (length s - (n + S l')) s)}
adamc@91 1127 + {forall l', S l' <= l
adamc@214 1128 -> ~ P (substring n (S l') s)
adamc@214 1129 \/ ~ P' (substring (n + S l') (length s - (n + S l')) s)}.
adamc@91 1130 refine (fix F (l : nat) : {exists l', S l' <= l
adamc@91 1131 /\ P (substring n (S l') s) /\ P' (substring (n + S l') (length s - (n + S l')) s)}
adamc@91 1132 + {forall l', S l' <= l
adamc@214 1133 -> ~ P (substring n (S l') s)
adamc@214 1134 \/ ~ P' (substring (n + S l') (length s - (n + S l')) s)} :=
adam@380 1135 match l return {exists l', S l' <= l
adam@380 1136 /\ P (substring n (S l') s) /\ P' (substring (n + S l') (length s - (n + S l')) s)}
adam@380 1137 + {forall l', S l' <= l
adam@380 1138 -> ~ P (substring n (S l') s)
adam@380 1139 \/ ~ P' (substring (n + S l') (length s - (n + S l')) s)} with
adamc@91 1140 | O => _
adamc@91 1141 | S l' =>
adamc@91 1142 (P_dec (substring n (S l') s) && P'_dec (n' := n + S l') _)
adamc@91 1143 || F l'
adamc@91 1144 end); clear F; crush; eauto 7;
adamc@91 1145 match goal with
adamc@91 1146 | [ H : ?X <= S ?Y |- _ ] => destruct (eq_nat_dec X (S Y)); crush
adamc@91 1147 end.
adamc@91 1148 Defined.
adamc@91 1149 End dec_star''.
adamc@91 1150
adamc@93 1151 (* begin hide *)
adamc@92 1152 Lemma star_length_contra : forall n,
adamc@92 1153 length s > n
adamc@92 1154 -> n >= length s
adamc@92 1155 -> False.
adamc@92 1156 crush.
adamc@92 1157 Qed.
adamc@92 1158
adamc@92 1159 Lemma star_length_flip : forall n n',
adamc@92 1160 length s - n <= S n'
adamc@92 1161 -> length s > n
adamc@92 1162 -> length s - n > 0.
adamc@92 1163 crush.
adamc@92 1164 Qed.
adamc@92 1165
adamc@92 1166 Hint Resolve star_length_contra star_length_flip substring_suffix_emp.
adamc@93 1167 (* end hide *)
adamc@92 1168
adamc@93 1169 (** The work of [dec_star''] is nested inside another linear search by [dec_star'], which provides the final functionality we need, but for arbitrary suffixes of [s], rather than just for [s] overall. *)
adamc@93 1170
adam@297 1171 Definition dec_star' : forall n n' : nat, length s - n' <= n
adamc@91 1172 -> {star P (substring n' (length s - n') s)}
adamc@214 1173 + {~ star P (substring n' (length s - n') s)}.
adamc@214 1174 refine (fix F (n n' : nat) : length s - n' <= n
adamc@91 1175 -> {star P (substring n' (length s - n') s)}
adamc@214 1176 + {~ star P (substring n' (length s - n') s)} :=
adamc@214 1177 match n with
adamc@91 1178 | O => fun _ => Yes
adamc@91 1179 | S n'' => fun _ =>
adamc@91 1180 le_gt_dec (length s) n'
adam@338 1181 || dec_star'' (n := n') (star P)
adam@338 1182 (fun n0 _ => Reduce (F n'' n0 _)) (length s - n')
adamc@92 1183 end); clear F; crush; eauto;
adamc@92 1184 match goal with
adamc@92 1185 | [ H : star _ _ |- _ ] => apply star_substring_inv in H; crush; eauto
adamc@92 1186 end;
adamc@92 1187 match goal with
adamc@92 1188 | [ H1 : _ < _ - _, H2 : forall l' : nat, _ <= _ - _ -> _ |- _ ] =>
adamc@92 1189 generalize (H2 _ (lt_le_S _ _ H1)); tauto
adamc@92 1190 end.
adamc@91 1191 Defined.
adamc@91 1192
adam@380 1193 (** Finally, we have [dec_star], defined by straightforward reduction from [dec_star']. *)
adamc@93 1194
adamc@214 1195 Definition dec_star : {star P s} + {~ star P s}.
adam@380 1196 refine (Reduce (dec_star' (n := length s) 0 _)); crush.
adamc@91 1197 Defined.
adamc@91 1198 End dec_star.
adamc@91 1199
adamc@93 1200 (* begin hide *)
adamc@86 1201 Lemma app_cong : forall x1 y1 x2 y2,
adamc@86 1202 x1 = x2
adamc@86 1203 -> y1 = y2
adamc@86 1204 -> x1 ++ y1 = x2 ++ y2.
adamc@86 1205 congruence.
adamc@86 1206 Qed.
adamc@86 1207
adamc@86 1208 Hint Resolve app_cong.
adamc@93 1209 (* end hide *)
adamc@93 1210
adamc@93 1211 (** With these helper functions completed, the implementation of our [matches] function is refreshingly straightforward. We only need one small piece of specific tactic work beyond what [crush] does for us. *)
adamc@86 1212
adam@297 1213 Definition matches : forall P (r : regexp P) s, {P s} + {~ P s}.
adamc@214 1214 refine (fix F P (r : regexp P) s : {P s} + {~ P s} :=
adamc@86 1215 match r with
adamc@86 1216 | Char ch => string_dec s (String ch "")
adamc@86 1217 | Concat _ _ r1 r2 => Reduce (split (F _ r1) (F _ r2) s)
adamc@87 1218 | Or _ _ r1 r2 => F _ r1 s || F _ r2 s
adamc@91 1219 | Star _ r => dec_star _ _ _
adamc@86 1220 end); crush;
adamc@86 1221 match goal with
adamc@86 1222 | [ H : _ |- _ ] => generalize (H _ _ (refl_equal _))
adamc@93 1223 end; tauto.
adamc@86 1224 Defined.
adamc@86 1225
adam@283 1226 (** It is interesting to pause briefly to consider alternate implementations of [matches]. Dependent types give us much latitude in how specific correctness properties may be encoded with types. For instance, we could have made [regexp] a non-indexed inductive type, along the lines of what is possible in traditional ML and Haskell. We could then have implemented a recursive function to map [regexp]s to their intended meanings, much as we have done with types and programs in other examples. That style is compatible with the [refine]-based approach that we have used here, and it might be an interesting exercise to redo the code from this subsection in that alternate style or some further encoding of the reader's choice. The main advantage of indexed inductive types is that they generally lead to the smallest amount of code. *)
adam@283 1227
adamc@93 1228 (* begin hide *)
adamc@86 1229 Example hi := Concat (Char "h"%char) (Char "i"%char).
adam@380 1230 Eval hnf in matches hi "hi".
adam@380 1231 Eval hnf in matches hi "bye".
adamc@87 1232
adamc@87 1233 Example a_b := Or (Char "a"%char) (Char "b"%char).
adam@380 1234 Eval hnf in matches a_b "".
adam@380 1235 Eval hnf in matches a_b "a".
adam@380 1236 Eval hnf in matches a_b "aa".
adam@380 1237 Eval hnf in matches a_b "b".
adam@283 1238 (* end hide *)
adam@283 1239
adam@380 1240 (** Many regular expression matching problems are easy to test. The reader may run each of the following queries to verify that it gives the correct answer. We use evaluation strategy %\index{tactics!hnf}%[hnf] to reduce each term to %\index{head-normal form}\emph{%#<i>#head-normal form#</i>#%}%, where the datatype constructor used to build its value is known. *)
adamc@91 1241
adamc@91 1242 Example a_star := Star (Char "a"%char).
adam@380 1243 Eval hnf in matches a_star "".
adam@380 1244 Eval hnf in matches a_star "a".
adam@380 1245 Eval hnf in matches a_star "b".
adam@380 1246 Eval hnf in matches a_star "aa".
adam@283 1247
adam@283 1248 (** Evaluation inside Coq does not scale very well, so it is easy to build other tests that run for hours or more. Such cases are better suited to execution with the extracted OCaml code. *)