author | Jean Pichon |
Thu, 29 Sep 2011 21:42:03 +0200 | |
changeset 45114 | fa3715b35370 |
parent 45015 | fdac1e9880eb |
child 46203 | d43ddad41d81 |
permissions | -rw-r--r-- |
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header{* Hoare Logic for Total Correctness *} |
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IMP/Util distinguishes between sets and functions again; imported only where used.
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theory HoareT imports Hoare_Sound_Complete begin |
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text{* |
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Now that we have termination, we can define |
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total validity, @{text"\<Turnstile>\<^sub>t"}, as partial validity and guaranteed termination:*} |
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definition hoare_tvalid :: "assn \<Rightarrow> com \<Rightarrow> assn \<Rightarrow> bool" |
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("\<Turnstile>\<^sub>t {(1_)}/ (_)/ {(1_)}" 50) where |
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"\<Turnstile>\<^sub>t {P}c{Q} \<equiv> \<forall>s. P s \<longrightarrow> (\<exists>t. (c,s) \<Rightarrow> t \<and> Q t)" |
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text{* Provability of Hoare triples in the proof system for total |
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correctness is written @{text"\<turnstile>\<^sub>t {P}c{Q}"} and defined |
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inductively. The rules for @{text"\<turnstile>\<^sub>t"} differ from those for |
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@{text"\<turnstile>"} only in the one place where nontermination can arise: the |
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@{term While}-rule. *} |
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inductive |
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hoaret :: "assn \<Rightarrow> com \<Rightarrow> assn \<Rightarrow> bool" ("\<turnstile>\<^sub>t ({(1_)}/ (_)/ {(1_)})" 50) |
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where |
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Skip: "\<turnstile>\<^sub>t {P} SKIP {P}" | |
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Assign: "\<turnstile>\<^sub>t {\<lambda>s. P(s[a/x])} x::=a {P}" | |
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Semi: "\<lbrakk> \<turnstile>\<^sub>t {P\<^isub>1} c\<^isub>1 {P\<^isub>2}; \<turnstile>\<^sub>t {P\<^isub>2} c\<^isub>2 {P\<^isub>3} \<rbrakk> \<Longrightarrow> \<turnstile>\<^sub>t {P\<^isub>1} c\<^isub>1;c\<^isub>2 {P\<^isub>3}" | |
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If: "\<lbrakk> \<turnstile>\<^sub>t {\<lambda>s. P s \<and> bval b s} c\<^isub>1 {Q}; \<turnstile>\<^sub>t {\<lambda>s. P s \<and> \<not> bval b s} c\<^isub>2 {Q} \<rbrakk> |
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\<Longrightarrow> \<turnstile>\<^sub>t {P} IF b THEN c\<^isub>1 ELSE c\<^isub>2 {Q}" | |
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While: |
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"\<lbrakk> \<And>n::nat. \<turnstile>\<^sub>t {\<lambda>s. P s \<and> bval b s \<and> f s = n} c {\<lambda>s. P s \<and> f s < n}\<rbrakk> |
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\<Longrightarrow> \<turnstile>\<^sub>t {P} WHILE b DO c {\<lambda>s. P s \<and> \<not>bval b s}" | |
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conseq: "\<lbrakk> \<forall>s. P' s \<longrightarrow> P s; \<turnstile>\<^sub>t {P}c{Q}; \<forall>s. Q s \<longrightarrow> Q' s \<rbrakk> \<Longrightarrow> |
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\<turnstile>\<^sub>t {P'}c{Q'}" |
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text{* The @{term While}-rule is like the one for partial correctness but it |
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requires additionally that with every execution of the loop body some measure |
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function @{term[source]"f :: state \<Rightarrow> nat"} decreases. *} |
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lemma strengthen_pre: |
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"\<lbrakk> \<forall>s. P' s \<longrightarrow> P s; \<turnstile>\<^sub>t {P} c {Q} \<rbrakk> \<Longrightarrow> \<turnstile>\<^sub>t {P'} c {Q}" |
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by (metis conseq) |
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lemma weaken_post: |
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"\<lbrakk> \<turnstile>\<^sub>t {P} c {Q}; \<forall>s. Q s \<longrightarrow> Q' s \<rbrakk> \<Longrightarrow> \<turnstile>\<^sub>t {P} c {Q'}" |
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by (metis conseq) |
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lemma Assign': "\<forall>s. P s \<longrightarrow> Q(s[a/x]) \<Longrightarrow> \<turnstile>\<^sub>t {P} x ::= a {Q}" |
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by (simp add: strengthen_pre[OF _ Assign]) |
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lemma While': |
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assumes "\<And>n::nat. \<turnstile>\<^sub>t {\<lambda>s. P s \<and> bval b s \<and> f s = n} c {\<lambda>s. P s \<and> f s < n}" |
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and "\<forall>s. P s \<and> \<not> bval b s \<longrightarrow> Q s" |
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shows "\<turnstile>\<^sub>t {P} WHILE b DO c {Q}" |
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by(blast intro: assms(1) weaken_post[OF While assms(2)]) |
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text{* Our standard example: *} |
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abbreviation "w n == |
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WHILE Less (V ''y'') (N n) |
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DO ( ''y'' ::= Plus (V ''y'') (N 1); ''x'' ::= Plus (V ''x'') (V ''y'') )" |
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lemma "\<turnstile>\<^sub>t {\<lambda>s. 0 <= n} ''x'' ::= N 0; ''y'' ::= N 0; w n {\<lambda>s. s ''x'' = \<Sum> {1 .. n}}" |
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apply(rule Semi) |
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prefer 2 |
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apply(rule While' |
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[where P = "\<lambda>s. s ''x'' = \<Sum> {1..s ''y''} \<and> 0 <= n \<and> 0 <= s ''y'' \<and> s ''y'' \<le> n" |
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and f = "\<lambda>s. nat n - nat (s ''y'')"]) |
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apply(rule Semi) |
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prefer 2 |
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apply(rule Assign) |
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apply(rule Assign') |
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apply (simp add: atLeastAtMostPlus1_int_conv algebra_simps) |
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apply clarsimp |
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apply arith |
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new fastforce replacing fastsimp - less confusing name
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apply fastforce |
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apply(rule Semi) |
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prefer 2 |
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apply(rule Assign) |
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apply(rule Assign') |
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apply simp |
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done |
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text{* The soundness theorem: *} |
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theorem hoaret_sound: "\<turnstile>\<^sub>t {P}c{Q} \<Longrightarrow> \<Turnstile>\<^sub>t {P}c{Q}" |
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proof(unfold hoare_tvalid_def, induct rule: hoaret.induct) |
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case (While P b f c) |
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show ?case |
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proof |
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fix s |
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show "P s \<longrightarrow> (\<exists>t. (WHILE b DO c, s) \<Rightarrow> t \<and> P t \<and> \<not> bval b t)" |
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proof(induction "f s" arbitrary: s rule: less_induct) |
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case (less n) |
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thus ?case by (metis While(2) WhileFalse WhileTrue) |
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qed |
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qed |
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next |
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case If thus ?case by auto blast |
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new fastforce replacing fastsimp - less confusing name
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qed fastforce+ |
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text{* |
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The completeness proof proceeds along the same lines as the one for partial |
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correctness. First we have to strengthen our notion of weakest precondition |
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to take termination into account: *} |
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definition wpt :: "com \<Rightarrow> assn \<Rightarrow> assn" ("wp\<^sub>t") where |
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"wp\<^sub>t c Q \<equiv> \<lambda>s. \<exists>t. (c,s) \<Rightarrow> t \<and> Q t" |
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lemma [simp]: "wp\<^sub>t SKIP Q = Q" |
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by(auto intro!: ext simp: wpt_def) |
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lemma [simp]: "wp\<^sub>t (x ::= e) Q = (\<lambda>s. Q(s(x := aval e s)))" |
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by(auto intro!: ext simp: wpt_def) |
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lemma [simp]: "wp\<^sub>t (c\<^isub>1;c\<^isub>2) Q = wp\<^sub>t c\<^isub>1 (wp\<^sub>t c\<^isub>2 Q)" |
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unfolding wpt_def |
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apply(rule ext) |
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apply auto |
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done |
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lemma [simp]: |
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"wp\<^sub>t (IF b THEN c\<^isub>1 ELSE c\<^isub>2) Q = (\<lambda>s. wp\<^sub>t (if bval b s then c\<^isub>1 else c\<^isub>2) Q s)" |
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apply(unfold wpt_def) |
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apply(rule ext) |
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apply auto |
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done |
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text{* Now we define the number of iterations @{term "WHILE b DO c"} needs to |
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terminate when started in state @{text s}. Because this is a truly partial |
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function, we define it as an (inductive) relation first: *} |
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inductive Its :: "bexp \<Rightarrow> com \<Rightarrow> state \<Rightarrow> nat \<Rightarrow> bool" where |
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Its_0: "\<not> bval b s \<Longrightarrow> Its b c s 0" | |
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Its_Suc: "\<lbrakk> bval b s; (c,s) \<Rightarrow> s'; Its b c s' n \<rbrakk> \<Longrightarrow> Its b c s (Suc n)" |
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text{* The relation is in fact a function: *} |
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lemma Its_fun: "Its b c s n \<Longrightarrow> Its b c s n' \<Longrightarrow> n=n'" |
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proof(induction arbitrary: n' rule:Its.induct) |
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(* new release: |
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case Its_0 thus ?case by(metis Its.cases) |
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next |
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case Its_Suc thus ?case by(metis Its.cases big_step_determ) |
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qed |
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*) |
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case Its_0 |
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from this(1) Its.cases[OF this(2)] show ?case by metis |
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next |
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case (Its_Suc b s c s' n n') |
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note C = this |
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from this(5) show ?case |
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proof cases |
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case Its_0 with Its_Suc(1) show ?thesis by blast |
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next |
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case Its_Suc with C show ?thesis by(metis big_step_determ) |
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qed |
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qed |
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text{* For all terminating loops, @{const Its} yields a result: *} |
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lemma WHILE_Its: "(WHILE b DO c,s) \<Rightarrow> t \<Longrightarrow> \<exists>n. Its b c s n" |
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proof(induction "WHILE b DO c" s t rule: big_step_induct) |
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case WhileFalse thus ?case by (metis Its_0) |
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next |
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case WhileTrue thus ?case by (metis Its_Suc) |
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qed |
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text{* Now the relation is turned into a function with the help of |
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the description operator @{text THE}: *} |
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definition its :: "bexp \<Rightarrow> com \<Rightarrow> state \<Rightarrow> nat" where |
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"its b c s = (THE n. Its b c s n)" |
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text{* The key property: every loop iteration increases @{const its} by 1. *} |
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lemma its_Suc: "\<lbrakk> bval b s; (c, s) \<Rightarrow> s'; (WHILE b DO c, s') \<Rightarrow> t\<rbrakk> |
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\<Longrightarrow> its b c s = Suc(its b c s')" |
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by (metis its_def WHILE_Its Its.intros(2) Its_fun the_equality) |
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lemma wpt_is_pre: "\<turnstile>\<^sub>t {wp\<^sub>t c Q} c {Q}" |
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proof (induction c arbitrary: Q) |
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case SKIP show ?case by simp (blast intro:hoaret.Skip) |
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next |
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case Assign show ?case by simp (blast intro:hoaret.Assign) |
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next |
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case Semi thus ?case by simp (blast intro:hoaret.Semi) |
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next |
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case If thus ?case by simp (blast intro:hoaret.If hoaret.conseq) |
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next |
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case (While b c) |
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let ?w = "WHILE b DO c" |
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{ fix n |
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have "\<forall>s. wp\<^sub>t ?w Q s \<and> bval b s \<and> its b c s = n \<longrightarrow> |
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wp\<^sub>t c (\<lambda>s'. wp\<^sub>t ?w Q s' \<and> its b c s' < n) s" |
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unfolding wpt_def by (metis WhileE its_Suc lessI) |
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note strengthen_pre[OF this While] |
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} note hoaret.While[OF this] |
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moreover have "\<forall>s. wp\<^sub>t ?w Q s \<and> \<not> bval b s \<longrightarrow> Q s" by (auto simp add:wpt_def) |
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ultimately show ?case by(rule weaken_post) |
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qed |
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text{*\noindent In the @{term While}-case, @{const its} provides the obvious |
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termination argument. |
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The actual completeness theorem follows directly, in the same manner |
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as for partial correctness: *} |
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theorem hoaret_complete: "\<Turnstile>\<^sub>t {P}c{Q} \<Longrightarrow> \<turnstile>\<^sub>t {P}c{Q}" |
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apply(rule strengthen_pre[OF _ wpt_is_pre]) |
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apply(auto simp: hoare_tvalid_def hoare_valid_def wpt_def) |
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done |
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end |