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Proofreading pass through Chapter 12

author | Adam Chlipala <adam@chlipala.net> |
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date | Tue, 28 Aug 2012 16:53:16 -0400 |

parents | 0d66f1a710b8 |

children | 1fd4109f7b31 |

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(* Copyright (c) 2009-2012, Adam Chlipala * * This work is licensed under a * Creative Commons Attribution-Noncommercial-No Derivative Works 3.0 * Unported License. * The license text is available at: * http://creativecommons.org/licenses/by-nc-nd/3.0/ *) (* begin hide *) Require Import List. Require Import DepList CpdtTactics. Set Implicit Arguments. (* end hide *) (** printing $ %({}*% #(<a/>*# *) (** printing ^ %*{})% #*<a/>)# *) (** %\chapter{Universes and Axioms}% *) (** 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. %\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. *) (** * The [Type] Hierarchy *) (** %\index{type hierarchy}%Every object in Gallina has a type. *) Check 0. (** %\vspace{-.15in}% [[ 0 : nat ]] It is natural enough that zero be considered as a natural number. *) Check nat. (** %\vspace{-.15in}% [[ nat : Set ]] From a set theory perspective, it is unsurprising to consider the natural numbers as a "set." *) Check Set. (** %\vspace{-.15in}% [[ Set : Type ]] 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]. *) Check Type. (** %\vspace{-.15in}% [[ Type : Type ]] 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? Let us repeat some of our queries after toggling a flag related to Coq's printing behavior.%\index{Vernacular commands!Set Printing Universes}% *) Set Printing Universes. Check nat. (** %\vspace{-.15in}% [[ nat : Set ]] *) Check Set. (** %\vspace{-.15in}% [[ Set : Type $ (0)+1 ^ ]] *) Check Type. (** %\vspace{-.15in}% [[ Type $ Top.3 ^ : Type $ (Top.3)+1 ^ ]] 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]. 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]. In the second 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. Another crucial concept in CIC is%\index{predicativity}% _predicativity_. Consider these queries. *) Check forall T : nat, fin T. (** %\vspace{-.15in}% [[ forall T : nat, fin T : Set ]] *) Check forall T : Set, T. (** %\vspace{-.15in}% [[ forall T : Set, T : Type $ max(0, (0)+1) ^ ]] *) Check forall T : Type, T. (** %\vspace{-.15in}% [[ forall T : Type $ Top.9 ^ , T : Type $ max(Top.9, (Top.9)+1) ^ ]] 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. 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]. *) Definition id (T : Set) (x : T) : T := x. Check id 0. (** %\vspace{-.15in}% [[ id 0 : nat Check id Set. ]] << Error: Illegal application (Type Error): ... The 1st term has type "Type (* (Top.15)+1 *)" which should be coercible to "Set". >> 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]. *) Reset id. Definition id (T : Type) (x : T) : T := x. Check id 0. (** %\vspace{-.15in}% [[ id 0 : nat ]] *) Check id Set. (** %\vspace{-.15in}% [[ id Set : Type $ Top.17 ^ ]] *) Check id Type. (** %\vspace{-.15in}% [[ id Type $ Top.18 ^ : Type $ Top.19 ^ ]] *) (** 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. [[ Check id id. ]] << Error: Universe inconsistency (cannot enforce Top.16 < Top.16). >> %\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, for any object of quantified type, none of those 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. *) (** ** Inductive Definitions *) (** 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 a reflected expression of type [T]. [[ Inductive exp : Set -> Set := | Const : forall T : Set, T -> exp T | Pair : forall T1 T2, exp T1 -> exp T2 -> exp (T1 * T2) | Eq : forall T, exp T -> exp T -> exp bool. ]] << Error: Large non-propositional inductive types must be in Type. >> 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. *) Inductive exp : Type -> Type := | Const : forall T, T -> exp T | Pair : forall T1 T2, exp T1 -> exp T2 -> exp (T1 * T2) | Eq : forall T, exp T -> exp T -> exp bool. (** 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]. That is the right behavior here, but it was wrong for the [Set] version of [exp]. Our new definition is accepted. We can build some sample expressions. *) Check Const 0. (** %\vspace{-.15in}% [[ Const 0 : exp nat ]] *) Check Pair (Const 0) (Const tt). (** %\vspace{-.15in}% [[ Pair (Const 0) (Const tt) : exp (nat * unit) ]] *) Check Eq (Const Set) (Const Type). (** %\vspace{-.15in}% [[ Eq (Const Set) (Const Type $ Top.59 ^ ) : exp bool ]] We can check many expressions, including fancy expressions that include types. However, it is not hard to hit a type-checking wall. [[ Check Const (Const O). ]] << Error: Universe inconsistency (cannot enforce Top.42 < Top.42). >> 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. *) (** [[ Print exp. ]] %\vspace{-.15in}%[[ Inductive exp : Type $ Top.8 ^ -> Type $ max(0, (Top.11)+1, (Top.14)+1, (Top.15)+1, (Top.19)+1) ^ := Const : forall T : Type $ Top.11 ^ , T -> exp T | Pair : forall (T1 : Type $ Top.14 ^ ) (T2 : Type $ Top.15 ^ ), exp T1 -> exp T2 -> exp (T1 * T2) | Eq : forall T : Type $ Top.19 ^ , exp T -> exp T -> exp bool ]] 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. A consequence of this is that [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. 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}% *) Print Universes. (** %\vspace{-.15in}% [[ Top.19 < Top.9 <= Top.8 Top.15 < Top.9 <= Top.8 <= Coq.Init.Datatypes.38 Top.14 < Top.9 <= Top.8 <= Coq.Init.Datatypes.37 Top.11 < Top.9 <= Top.8 ]] 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. 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. *) (* begin hide *) (* begin thide *) Inductive prod := pair. Reset prod. (* end thide *) (* end hide *) (** %\vspace{-.3in}%[[ Print prod. ]] %\vspace{-.15in}%[[ Inductive prod (A : Type $ Coq.Init.Datatypes.37 ^ ) (B : Type $ Coq.Init.Datatypes.38 ^ ) : Type $ max(Coq.Init.Datatypes.37, Coq.Init.Datatypes.38) ^ := pair : A -> B -> A * B ]] 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]. 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. %\medskip% Something interesting is revealed in the annotated definition of [prod]. 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. 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. *) Check (nat, (Type, Set)). (** %\vspace{-.15in}% [[ (nat, (Type $ Top.44 ^ , Set)) : Set * (Type $ Top.45 ^ * Type $ Top.46 ^ ) ]] The same cannot be done with a counterpart to [prod] that does not use parameters. *) Inductive prod' : Type -> Type -> Type := | pair' : forall A B : Type, A -> B -> prod' A B. (** %\vspace{-.15in}%[[ Check (pair' nat (pair' Type Set)). ]] << Error: Universe inconsistency (cannot enforce Top.51 < Top.51). >> 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. 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. *) Inductive foo (A : Type) : Type := | Foo : A -> foo A. (* begin hide *) Unset Printing Universes. (* end hide *) Check foo nat. (** %\vspace{-.15in}% [[ foo nat : Set ]] *) Check foo Set. (** %\vspace{-.15in}% [[ foo Set : Type ]] *) Check foo True. (** %\vspace{-.15in}% [[ foo True : Prop ]] 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]. Imitation polymorphism can be confusing in some contexts. For instance, it is what is responsible for this weird behavior. *) Inductive bar : Type := Bar : bar. Check bar. (** %\vspace{-.15in}% [[ bar : Prop ]] The type that Coq comes up with may be used in strictly more contexts than the type one might have expected. *) (** ** Deciphering Baffling Messages About Inability to Unify *) (** 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. *) Theorem symmetry : forall A B : Type, A = B -> B = A. intros ? ? H; rewrite H; reflexivity. Qed. (** Let us attempt an admittedly silly proof of the following theorem. *) Theorem illustrative_but_silly_detour : unit = unit. (** %\vspace{-.25in}%[[ apply symmetry. ]] << Error: Impossible to unify "?35 = ?34" with "unit = unit". >> 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! The following command is the secret to getting better error messages in such cases: *) Set Printing All. (** %\vspace{-.15in}%[[ apply symmetry. ]] << Error: Impossible to unify "@eq Type ?46 ?45" with "@eq Set unit unit". >> 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. *) Abort. (** 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. *) Theorem illustrative_but_silly_detour : (unit : Type) = unit. apply symmetry; reflexivity. Qed. (** 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. 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: *) Unset Printing All. Theorem ex_symmetry : (exists x, x = 0) -> (exists x, 0 = x). eexists. (** %\vspace{-.15in}%[[ H : exists x : nat, x = 0 ============================ 0 = ?98 ]] *) destruct H. (** %\vspace{-.15in}%[[ x : nat H : x = 0 ============================ 0 = ?99 ]] *) (** %\vspace{-.2in}%[[ symmetry; exact H. ]] << Error: In environment x : nat H : x = 0 The term "H" has type "x = 0" while it is expected to have type "?99 = 0". >> 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. *) Restart. destruct 1 as [x]; apply ex_intro with x; symmetry; assumption. Qed. (** 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]. *) (** * The [Prop] Universe *) (** 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. 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]. *) Print sig. (** %\vspace{-.15in}% [[ Inductive sig (A : Type) (P : A -> Prop) : Type := exist : forall x : A, P x -> sig P ]] *) Print ex. (** %\vspace{-.15in}% [[ Inductive ex (A : Type) (P : A -> Prop) : Prop := ex_intro : forall x : A, P x -> ex P ]] It is natural to want a function to extract the first components of data structures like these. Doing so is easy enough for [sig]. *) Definition projS A (P : A -> Prop) (x : sig P) : A := match x with | exist v _ => v end. (* begin hide *) (* begin thide *) Definition projE := O. (* end thide *) (* end hide *) (** We run into trouble with a version that has been changed to work with [ex]. [[ Definition projE A (P : A -> Prop) (x : ex P) : A := match x with | ex_intro v _ => v end. ]] << Error: Incorrect elimination of "x" in the inductive type "ex": the return type has sort "Type" while it should be "Prop". Elimination of an inductive object of sort Prop is not allowed on a predicate in sort Type because proofs can be eliminated only to build proofs. >> 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. 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. 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. *) Definition sym_sig (x : sig (fun n => n = 0)) : sig (fun n => 0 = n) := match x with | exist n pf => exist _ n (sym_eq pf) end. Extraction sym_sig. (** << (** val sym_sig : nat -> nat **) let sym_sig x = x >> 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. *) Definition sym_ex (x : ex (fun n => n = 0)) : ex (fun n => 0 = n) := match x with | ex_intro n pf => ex_intro _ n (sym_eq pf) end. Extraction sym_ex. (** << (** val sym_ex : __ **) let sym_ex = __ >> 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. 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. 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. %\medskip% 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. *) Check forall P Q : Prop, P \/ Q -> Q \/ P. (** %\vspace{-.15in}% [[ forall P Q : Prop, P \/ Q -> Q \/ P : Prop ]] 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. Impredicativity also allows us to implement a version of our earlier [exp] type that does not suffer from the weakness that we found. *) Inductive expP : Type -> Prop := | ConstP : forall T, T -> expP T | PairP : forall T1 T2, expP T1 -> expP T2 -> expP (T1 * T2) | EqP : forall T, expP T -> expP T -> expP bool. Check ConstP 0. (** %\vspace{-.15in}% [[ ConstP 0 : expP nat ]] *) Check PairP (ConstP 0) (ConstP tt). (** %\vspace{-.15in}% [[ PairP (ConstP 0) (ConstP tt) : expP (nat * unit) ]] *) Check EqP (ConstP Set) (ConstP Type). (** %\vspace{-.15in}% [[ EqP (ConstP Set) (ConstP Type) : expP bool ]] *) Check ConstP (ConstP O). (** %\vspace{-.15in}% [[ ConstP (ConstP 0) : expP (expP nat) ]] 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 [=]. *) Inductive eqPlus : forall T, T -> T -> Prop := | Base : forall T (x : T), eqPlus x x | Func : forall dom ran (f1 f2 : dom -> ran), (forall x : dom, eqPlus (f1 x) (f2 x)) -> eqPlus f1 f2. Check (Base 0). (** %\vspace{-.15in}% [[ Base 0 : eqPlus 0 0 ]] *) Check (Func (fun n => n) (fun n => 0 + n) (fun n => Base n)). (** %\vspace{-.15in}% [[ Func (fun n : nat => n) (fun n : nat => 0 + n) (fun n : nat => Base n) : eqPlus (fun n : nat => n) (fun n : nat => 0 + n) ]] *) Check (Base (Base 1)). (** %\vspace{-.15in}% [[ Base (Base 1) : eqPlus (Base 1) (Base 1) ]] *) (** Stating equality facts about proofs may seem baroque, but we have already seen its utility in the chapter on reasoning about equality proofs. *) (** * Axioms *) (** 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. 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. *) (** ** The Basics *) (** One simple example of a useful axiom is the %\index{law of the excluded middle}%law of the excluded middle. *) Require Import Classical_Prop. Print classic. (** %\vspace{-.15in}% [[ *** [ classic : forall P : Prop, P \/ ~ P ] ]] In the implementation of module [Classical_Prop], this axiom was defined with the command%\index{Vernacular commands!Axiom}% *) Axiom classic : forall P : Prop, P \/ ~ P. (** 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. *) Parameter num : nat. Axiom positive : num > 0. Reset num. (** 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. 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]. *) Axiom not_classic : ~ forall P : Prop, P \/ ~ P. Theorem uhoh : False. generalize classic not_classic; tauto. Qed. Theorem uhoh_again : 1 + 1 = 3. destruct uhoh. Qed. Reset not_classic. (** 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. Recall that Coq implements%\index{constructive logic}% _constructive_ logic by default, where 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. 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. 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}% *) Theorem t1 : forall P : Prop, P -> ~ ~ P. tauto. Qed. Print Assumptions t1. (** << Closed under the global context >> *) Theorem t2 : forall P : Prop, ~ ~ P -> P. (** %\vspace{-.25in}%[[ tauto. ]] << Error: tauto failed. >> *) intro P; destruct (classic P); tauto. Qed. Print Assumptions t2. (** %\vspace{-.15in}% [[ Axioms: classic : forall P : Prop, P \/ ~ P ]] It is possible to avoid this dependence in some specific cases, where excluded middle _is_ provable, for decidable families of propositions. *) Theorem nat_eq_dec : forall n m : nat, n = m \/ n <> m. induction n; destruct m; intuition; generalize (IHn m); intuition. Qed. Theorem t2' : forall n m : nat, ~ ~ (n = m) -> n = m. intros n m; destruct (nat_eq_dec n m); tauto. Qed. Print Assumptions t2'. (** << Closed under the global context >> %\bigskip% 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. *) Require Import ProofIrrelevance. Print proof_irrelevance. (** %\vspace{-.15in}% [[ *** [ proof_irrelevance : forall (P : Prop) (p1 p2 : P), p1 = p2 ] ]] This axiom asserts that any two proofs of the same proposition are equal. Recall this example function from Chapter 6. *) (* begin hide *) Lemma zgtz : 0 > 0 -> False. crush. Qed. (* end hide *) Definition pred_strong1 (n : nat) : n > 0 -> nat := match n with | O => fun pf : 0 > 0 => match zgtz pf with end | S n' => fun _ => n' end. (** We might want to prove that different proofs of [n > 0] do not lead to different results from our richly typed predecessor function. *) Theorem pred_strong1_irrel : forall n (pf1 pf2 : n > 0), pred_strong1 pf1 = pred_strong1 pf2. destruct n; crush. Qed. (** 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]. *) Theorem pred_strong1_irrel' : forall n (pf1 pf2 : n > 0), pred_strong1 pf1 = pred_strong1 pf2. intros; f_equal; apply proof_irrelevance. Qed. (** %\bigskip% 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: *) Require Import Eqdep. Import Eq_rect_eq. Print eq_rect_eq. (** %\vspace{-.15in}% [[ *** [ eq_rect_eq : forall (U : Type) (p : U) (Q : U -> Type) (x : Q p) (h : p = p), x = eq_rect p Q x p h ] ]] 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." *) Corollary UIP_refl : forall A (x : A) (pf : x = x), pf = eq_refl x. intros; replace pf with (eq_rect x (eq x) (eq_refl x) x pf); [ symmetry; apply eq_rect_eq | exact (match pf as pf' return match pf' in _ = y return x = y with | eq_refl => eq_refl x end = pf' with | eq_refl => eq_refl _ end) ]. Qed. Corollary UIP : forall A (x y : A) (pf1 pf2 : x = y), pf1 = pf2. intros; generalize pf1 pf2; subst; intros; match goal with | [ |- ?pf1 = ?pf2 ] => rewrite (UIP_refl pf1); rewrite (UIP_refl pf2); reflexivity end. Qed. (* begin hide *) (* begin thide *) Require Eqdep_dec. (* end thide *) (* end hide *) (** 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. 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. %\bigskip% There are two more basic axioms that are often assumed, to avoid complications that do not arise in set theory. *) Require Import FunctionalExtensionality. Print functional_extensionality_dep. (** %\vspace{-.15in}% [[ *** [ functional_extensionality_dep : forall (A : Type) (B : A -> Type) (f g : forall x : A, B x), (forall x : A, f x = g x) -> f = g ] ]] 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. A simple corollary shows that the same property applies to predicates. *) Corollary predicate_extensionality : forall (A : Type) (B : A -> Prop) (f g : forall x : A, B x), (forall x : A, f x = g x) -> f = g. intros; apply functional_extensionality_dep; assumption. Qed. (** In some cases, one might prefer to assert this corollary as the axiom, to restrict the consequences to proofs and not programs. *) (** ** Axioms of Choice *) (** 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. First, it is possible to implement a choice operator _without_ axioms in some potentially surprising cases. *) Require Import ConstructiveEpsilon. Check constructive_definite_description. (** %\vspace{-.15in}% [[ constructive_definite_description : forall (A : Set) (f : A -> nat) (g : nat -> A), (forall x : A, g (f x) = x) -> forall P : A -> Prop, (forall x : A, {P x} + {~ P x}) -> (exists! x : A, P x) -> {x : A | P x} ]] *) Print Assumptions constructive_definite_description. (** << Closed under the global context >> 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. 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. *) Require Import ClassicalUniqueChoice. Check dependent_unique_choice. (** %\vspace{-.15in}% [[ dependent_unique_choice : forall (A : Type) (B : A -> Type) (R : forall x : A, B x -> Prop), (forall x : A, exists! y : B x, R x y) -> exists f : forall x : A, B x, forall x : A, R x (f x) ]] 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. *) (* begin hide *) (* begin thide *) Require RelationalChoice. (* end thide *) (* end hide *) Require Import ClassicalChoice. Check choice. (** %\vspace{-.15in}% [[ choice : forall (A B : Type) (R : A -> B -> Prop), (forall x : A, exists y : B, R x y) -> exists f : A -> B, forall x : A, R x (f x) ]] This principle is proved as a theorem, based on the unique choice axiom and an additional axiom of relational choice from the [RelationalChoice] module. 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. *) Definition choice_Set (A B : Type) (R : A -> B -> Prop) (H : forall x : A, {y : B | R x y}) : {f : A -> B | forall x : A, R x (f x)} := exist (fun f => forall x : A, R x (f x)) (fun x => proj1_sig (H x)) (fun x => proj2_sig (H x)). (** %\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. 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 truly is more than repackaging a function with a different type. %\bigskip% 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 contrasts with [Prop], where the restriction applies even to non-large inductive types, and where the result type may only have type [Prop]. 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, this can lead to inconsistency. Impredicative [Set] can be useful for modeling inherently impredicative mathematical concepts, but almost all Coq developments get by fine without it. *) (** ** Axioms and Computation *) (** 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. *) Definition cast (x y : Set) (pf : x = y) (v : x) : y := match pf with | eq_refl => v end. (** Computation over programs that use [cast] can proceed smoothly. *) Eval compute in (cast (eq_refl (nat -> nat)) (fun n => S n)) 12. (** %\vspace{-.15in}%[[ = 13 : nat ]] *) (** Things do not go as smoothly when we use [cast] with proofs that rely on axioms. *) Theorem t3 : (forall n : nat, fin (S n)) = (forall n : nat, fin (n + 1)). change ((forall n : nat, (fun n => fin (S n)) n) = (forall n : nat, (fun n => fin (n + 1)) n)); rewrite (functional_extensionality (fun n => fin (n + 1)) (fun n => fin (S n))); crush. Qed. Eval compute in (cast t3 (fun _ => First)) 12. (** %\vspace{-.15in}%[[ = match t3 in (_ = P) return P with | eq_refl => fun n : nat => First end 12 : fin (12 + 1) ]] 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. *) Reset t3. Theorem t3 : (forall n : nat, fin (S n)) = (forall n : nat, fin (n + 1)). change ((forall n : nat, (fun n => fin (S n)) n) = (forall n : nat, (fun n => fin (n + 1)) n)); rewrite (functional_extensionality (fun n => fin (n + 1)) (fun n => fin (S n))); crush. Defined. Eval compute in (cast t3 (fun _ => First)) 12. (** %\vspace{-.15in}%[[ = match match match functional_extensionality .... ]] 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. If we are careful in using tactics to prove an equality, we can still compute with casts over the proof. *) Lemma plus1 : forall n, S n = n + 1. induction n; simpl; intuition. Defined. Theorem t4 : forall n, fin (S n) = fin (n + 1). intro; f_equal; apply plus1. Defined. Eval compute in cast (t4 13) First. (** %\vspace{-.15in}% [[ = First : fin (13 + 1) ]] 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. *) (** ** Methods for Avoiding Axioms *) (** 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. 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!) 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: *) Theorem fin_cases : forall n (f : fin (S n)), f = First \/ exists f', f = Next f'. intros; dep_destruct f; eauto. Qed. (* begin hide *) Require Import JMeq. (* begin thide *) Definition jme := (JMeq, JMeq_eq). (* end thide *) (* end hide *) Print Assumptions fin_cases. (** %\vspace{-.15in}%[[ Axioms: JMeq_eq : forall (A : Type) (x y : A), JMeq x y -> x = y ]] 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. *) (* begin thide *) Lemma fin_cases_again' : forall n (f : fin n), match n return fin n -> Prop with | O => fun _ => False | S n' => fun f => f = First \/ exists f', f = Next f' end f. destruct f; eauto. Qed. (** 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. *) Theorem fin_cases_again : forall n (f : fin (S n)), f = First \/ exists f', f = Next f'. intros; exact (fin_cases_again' f). Qed. (* end thide *) Print Assumptions fin_cases_again. (** %\vspace{-.15in}% << Closed under the global context >> *) (* begin thide *) (** 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. 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. *) Definition finOut n (f : fin n) : match n return fin n -> Type with | O => fun _ => Empty_set | _ => fun f => {f' : _ | f = Next f'} + {f = First} end f := match f with | First _ => inright _ (eq_refl _) | Next _ f' => inleft _ (exist _ f' (eq_refl _)) end. (* end thide *) (** 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!) *) Inductive formula : list Type -> Type := | Inject : forall Ts, Prop -> formula Ts | VarEq : forall T Ts, T -> formula (T :: Ts) | Lift : forall T Ts, formula Ts -> formula (T :: Ts) | Forall : forall T Ts, formula (T :: Ts) -> formula Ts | And : forall Ts, formula Ts -> formula Ts -> formula Ts. (** 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. *) Inductive proof : formula nil -> Prop := | PInject : forall (P : Prop), P -> proof (Inject nil P) | PAnd : forall p q, proof p -> proof q -> proof (And p q). (** Let us prove a lemma showing that a "[P /\ Q -> P]" rule is derivable within the rules of [proof]. *) Theorem proj1 : forall p q, proof (And p q) -> proof p. destruct 1. (** %\vspace{-.15in}%[[ p : formula nil q : formula nil P : Prop H : P ============================ proof p ]] *) (** 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].) 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. *) Restart. Require Import Program. intros ? ? H; dependent destruction H; auto. Qed. Print Assumptions proj1. (** %\vspace{-.15in}%[[ Axioms: eq_rect_eq : forall (U : Type) (p : U) (Q : U -> Type) (x : Q p) (h : p = p), x = eq_rect p Q x p h ]] 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. *) Lemma proj1_again' : forall r, proof r -> forall p q, r = And p q -> proof p. destruct 1; crush. (** %\vspace{-.15in}%[[ H0 : Inject [] P = And p q ============================ proof p ]] The first goal looks reasonable. Hypothesis [H0] is clearly contradictory, as [discriminate] can show. *) discriminate. (** %\vspace{-.15in}%[[ H : proof p H1 : And p q = And p0 q0 ============================ proof p0 ]] It looks like we are almost done. Hypothesis [H1] gives [p = p0] by injectivity of constructors, and then [H] finishes the case. *) injection H1; intros. (* begin hide *) (* begin thide *) Definition existT' := existT. (* end thide *) (* end hide *) (** Unfortunately, the "equality" that we expected between [p] and [p0] comes in a strange form: [[ H3 : existT (fun Ts : list Type => formula Ts) []%list p = existT (fun Ts : list Type => formula Ts) []%list p0 ============================ proof p0 ]] 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. How exactly does an axiom come into the picture here? Let us ask [crush] to finish the proof. *) crush. Qed. Print Assumptions proj1_again'. (** %\vspace{-.15in}%[[ Axioms: eq_rect_eq : forall (U : Type) (p : U) (Q : U -> Type) (x : Q p) (h : p = p), x = eq_rect p Q x p h ]] 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. Hope is not lost, however. We can produce an even stranger looking lemma, which gives us the theorem without axioms. *) Lemma proj1_again'' : forall r, proof r -> match r with | And Ps p _ => match Ps return formula Ps -> Prop with | nil => fun p => proof p | _ => fun _ => True end p | _ => True end. destruct 1; auto. Qed. Theorem proj1_again : forall p q, proof (And p q) -> proof p. intros ? ? H; exact (proj1_again'' H). Qed. Print Assumptions proj1_again. (** << Closed under the global context >> This example illustrates again how some of the same design patterns we learned for dependently typed programming can be used fruitfully in theorem statements. %\medskip% 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. 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. *) Section withTypes. Variable types : list Set. (** To give the semantics of terms, we will need to represent value environments, which assign each variable a term of the proper type. *) Variable values : hlist (fun x : Set => x) types. (** 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. *) Variable natIndex : nat. Variable natIndex_ok : nth_error types natIndex = Some nat. (** 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. *) Lemma nth_error_nil : forall A n x, nth_error (@nil A) n = Some x -> False. destruct n; simpl; unfold error; congruence. Defined. Implicit Arguments nth_error_nil [A n x]. Lemma Some_inj : forall A (x y : A), Some x = Some y -> x = y. congruence. Defined. Fixpoint getNat (types' : list Set) (values' : hlist (fun x : Set => x) types') (natIndex : nat) : (nth_error types' natIndex = Some nat) -> nat := match values' with | HNil => fun pf => match nth_error_nil pf with end | HCons t ts x values'' => match natIndex return nth_error (t :: ts) natIndex = Some nat -> nat with | O => fun pf => match Some_inj pf in _ = T return T with | eq_refl => x end | S natIndex' => getNat values'' natIndex' end end. End withTypes. (** The problem becomes apparent when we experiment with running [getNat] on a concrete [types] list. *) Definition myTypes := unit :: nat :: bool :: nil. Definition myValues : hlist (fun x : Set => x) myTypes := tt ::: 3 ::: false ::: HNil. Definition myNatIndex := 1. Theorem myNatIndex_ok : nth_error myTypes myNatIndex = Some nat. reflexivity. Defined. Eval compute in getNat myValues myNatIndex myNatIndex_ok. (** %\vspace{-.15in}%[[ = 3 ]] 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. *) Theorem getNat_is_reasonable : forall pf, getNat myValues myNatIndex pf = 3. intro; compute. (** << 1 subgoal >> %\vspace{-.3in}%[[ pf : nth_error myTypes myNatIndex = Some nat ============================ match match pf in (_ = y) return (nat = match y with | Some H => H | None => nat end) with | eq_refl => eq_refl end in (_ = T) return T with | eq_refl => 3 end = 3 ]] 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. *) Abort. (** 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. *) Fixpoint copies A (x : A) (n : nat) : list A := match n with | O => nil | S n' => x :: copies x n' end. Fixpoint update A (ls : list A) (n : nat) (x : A) : list A := match ls with | nil => copies x n ++ x :: nil | y :: ls' => match n with | O => x :: ls' | S n' => y :: update ls' n' x end end. (** Now let us revisit the definition of [getNat]. *) Section withTypes'. Variable types : list Set. Variable natIndex : nat. (** 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. *) Definition types' := update types natIndex nat. Variable values : hlist (fun x : Set => x) types'. (** Now a bit of dependent pattern matching helps us rewrite [getNat] in a way that avoids any use of equality proofs. *) Fixpoint skipCopies (n : nat) : hlist (fun x : Set => x) (copies nat n ++ nat :: nil) -> nat := match n with | O => fun vs => hhd vs | S n' => fun vs => skipCopies n' (htl vs) end. Fixpoint getNat' (types'' : list Set) (natIndex : nat) : hlist (fun x : Set => x) (update types'' natIndex nat) -> nat := match types'' with | nil => skipCopies natIndex | t :: types0 => match natIndex return hlist (fun x : Set => x) (update (t :: types0) natIndex nat) -> nat with | O => fun vs => hhd vs | S natIndex' => fun vs => getNat' types0 natIndex' (htl vs) end end. End withTypes'. (** 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. *) Theorem getNat_is_reasonable : getNat' myTypes myNatIndex myValues = 3. reflexivity. Qed. (** The same parameters as before work without alteration, and we avoid use of axioms. *)