hol88-contrib-source 2.02.19940316-19 / usr / share / hol88-2.02.19940316 / contrib / rule-induction / opsem.ml

% ===================================================================== %
% FILE		: opsem.ml						%
% DESCRIPTION   : Creates a theory of the syntax and operational 	%
% 		  semantics of a very simple imperative programming 	%
% 		  language.  Illustrates the inductive definitions 	%
% 		  package with proofs that the evaluation relation for	%
%		  the given semantics is deterministic and that the 	%
%		  Hoare-logic rule for while loops follows from a 	%
%		  suitable definition of partial correctness.		%
%									%
% AUTHORS	: Tom Melham and Juanito Camilleri			%
% DATE		: 91.10.09						%
% ===================================================================== %

% --------------------------------------------------------------------- %
% Open a new theory and load the required libraries.			%
% --------------------------------------------------------------------- %

new_theory `opsem`;;

load_library `string`;;

load_library `ind_defs`;;

% ===================================================================== %
% Syntax of the programming language.					%
% ===================================================================== %

% --------------------------------------------------------------------- %
% Program variables will be represented by strings, and states will be  %
% modelled by functions from program variables to natural numbers.	%
% --------------------------------------------------------------------- %

new_type_abbrev (`state`, ":string->num");;

% --------------------------------------------------------------------- %
% Natural number expressions and boolean expressions will just be 	%
% modelled by total functions from states to numbers and booleans       %
% respectively.  This simplification allows us to concentrate in this   %
% example on defining the semantics of commands.			%
% --------------------------------------------------------------------- %

new_type_abbrev(`nexp`, ":state->num");;
new_type_abbrev(`bexp`, ":state->bool");;

% --------------------------------------------------------------------- %
% We can now use the recursive types package to define the syntax of    %
% commands (or `programs').  We have the following types of commands:   %
%									%
%    C ::=   skip                     (does nothing)			%
%          | V := E                   (assignment)          		%
%          | C1 ; C2                  (sequencing)           		%
%          | if B then C1 else C2     (conditional)			%
%          | while B do C             (repetition)			%
%                                                                       %
% where V ranges over program varibles, E ranges over natural number	%
% expressions, B ranges over boolean expressions, and C, C1 and C2      %
% range over commands.  						%
%									%
% In the logic, we represent this abstract syntax with a set of prefix  %
% type constructors. So we have:					%
%						                        %
%    V := E                  represented by "assign V E"		%
%    C1 ; C2                 represented by "seq C1 C2"			%
%    if B then C1 else C2    represented by "if B C1 C2"		%
%    while B do C            represented by "while B C"			%
%									%
% For notational clarity, we later introduce two constants := and ; as 	%
% infix abbreviations for assign and seq.  This can't be done here just %
% because define_type doesn't suppport infix constructors.  		%
% --------------------------------------------------------------------- %

let comm = 
    define_type `comm`
    `comm =   skip
            | assign string nexp
	    | seq comm comm        
            | if bexp comm comm    
	    | while bexp comm`;;


% --------------------------------------------------------------------- %
% Define an infix function `:=' for assignment.				%
% --------------------------------------------------------------------- %

let assign_def = 
    new_infix_definition 
    (`assign_def`,"$:= V E = assign V E");;


% --------------------------------------------------------------------- %
% Define infix function `;' for sequencing.				%
% --------------------------------------------------------------------- %

let seq_def = 
    new_infix_definition 
    (`seq_def`,"$; C1 C2 = seq C1 C2");;


% --------------------------------------------------------------------- %
% Replace seq and assign by the infixes := and ;.			%
% --------------------------------------------------------------------- %

let comm = 
     save_thm
      (`comm_thm`,
        PURE_ONCE_REWRITE_RULE
	  [SYM (SPEC_ALL assign_def);SYM (SPEC_ALL seq_def)]
          comm);;

% ===================================================================== %
% Standard syntactic theory, derived by the recursive types package.	%
% ===================================================================== %

% --------------------------------------------------------------------- %
% Structural induction theorem for commands.				%
% --------------------------------------------------------------------- %

let induct = 
    save_thm (`induct`,prove_induction_thm comm);;

% --------------------------------------------------------------------- %
% Exhaustive case analysis theorem for commands.			%
% --------------------------------------------------------------------- %

let cases = 
    save_thm (`cases`, prove_cases_thm induct);;

% --------------------------------------------------------------------- %
% Prove that the abstract syntax constructors are one-to-one.		%
% --------------------------------------------------------------------- %

let const11 = 
    let [assign11;seq11;if11;while11] =
        (CONJUNCTS (prove_constructors_one_one comm)) in
        map save_thm
        [(`assign11`,assign11);
         (`seq11`,seq11);
         (`if11`,if11);
         (`while11`,while11)];;


% --------------------------------------------------------------------- %
% Prove that the constructors yield syntactically distinct values. Note %
% that one typically needs symmetric forms of the inequalities, so 	%
% these	are constructed here and grouped togther into one theorem.	%
% --------------------------------------------------------------------- %

let distinct =
    let ths = CONJUNCTS (prove_constructors_distinct comm) in
    let rths = map (GEN_ALL o NOT_EQ_SYM o SPEC_ALL) ths in
        save_thm(`distinct`, LIST_CONJ (ths @ rths));;

% ===================================================================== %
% Definition of the operational semantics.				%
% ===================================================================== %

% --------------------------------------------------------------------- %
% The semantics of commands will be given by an evaluation relation	%
%									%
%       EVAL : comm -> state -> state -> bool     			%
%									%
% defined inductively such that 					%
%									%
%       EVAL C s1 s2							%
%									%
% holds exactly when executing the command C in the initial state s1 	%
% terminates in the final state s2. The evaluation relation is defined  %
% inductively by the set of rules shown below.  			%
% --------------------------------------------------------------------- %

let rules,ind = 
    let EVAL = "EVAL : comm -> state -> state -> bool" in
    new_inductive_definition false `trans` 

    ("^EVAL C s1 s2", []) 
    
     [ [
       % ------------------------------------------------- % ],
                         "^EVAL skip s s"                  ;  

       [              
       % ------------------------------------------------- % ],
            "^EVAL (V := E) s (\v. (v=V) => E s | s v)"   ;


       [    "^EVAL C1 s1 s2";        "^EVAL C2 s2 s3"     
       % ------------------------------------------------- % ],
                      "^EVAL (C1;C2) s1 s3"               ;


       [               "^EVAL C1 s1 s2"                   ;
       % ------------------------------------------------- % "(B:bexp) s1"],
                   "^EVAL (if B C1 C2) s1 s2"             ;


       [                 "^EVAL C2 s1 s2"                 ;
       % ------------------------------------------------- % "~((B:bexp) s1)"],
                    "^EVAL (if B C1 C2) s1 s2"            ;

       [                                                  
       % ------------------------------------------------- % "~((B:bexp) s)"],
                     "^EVAL (while B C) s s"              ;


       [   "^EVAL C s1 s2";    "^EVAL (while B C) s2 s3"   ;
       % ------------------------------------------------- %  "(B:bexp) s1"],
                     "^EVAL (while B C) s1 s3"             ];;

% --------------------------------------------------------------------- %
% Stronger form of rule induction.					%
% --------------------------------------------------------------------- %

let sind = derive_strong_induction(rules,ind);;

% --------------------------------------------------------------------- %
% Construct the standard rule induction tactic for EVAL.  This uses	%
% the `weaker' form of the rule induction theorem, and both premisses	%
% and side conditions are simply assumed (in stripped form).  This 	%
% served for many proofs, but when a more elaborate treatment of	%
% premisses or side conditions is needed, or when the stronger form of	%
% induction is required, a specialized  rule induction tactic is	%
% constructed on the fly.						%
% --------------------------------------------------------------------- %

let RULE_INDUCT_TAC =
    RULE_INDUCT_THEN ind STRIP_ASSUME_TAC STRIP_ASSUME_TAC;;

% --------------------------------------------------------------------- %
% Prove the case analysis theorem for the evaluation rules.		%
% --------------------------------------------------------------------- %

let ecases = derive_cases_thm (rules,ind);;

% ===================================================================== %
% Derivation of backwards case analysis theorems for each rule.         %
%									%
% These theorems are consequences of the general case analysis theorem  %
% proved above.  They are used to justify formal reasoning in which the %
% rules are driven `backwards', inferring premisses from conclusions.   %
% One infers from the assertion that:					%
%									%
%    1: EVAL C s1 s2 							%
%									%
% for a specific command C (e.g. for C = "skip") that the corresponding %
% instance of the premisses of the rule(s) for C must also hold, since  %
% (1) can hold only by virtue of being derivable by the rule for C.     %
% This kind of reasoning occurs frequently in proofs about operational  %
% semantics.  Formally, one must use the fact that the logical 		%
% representations of syntactically different commands are distinct, a   %
% fact automatically proved by the recursive types package.		%
% ===================================================================== %

% --------------------------------------------------------------------- %
% The following rule is used to simplify special cases of the general   %
% exhaustive case analysis theorem, which looks something like:		%
%									%
%      |- !C s1 s2.							%
%          EVAL C s1 s2 =						%
%             (C = skip) ... \/						%
%             (?V E. (C = V := E) ...) \/				%
%             (?C1 C2 s2'. (C = C1 ; C2) ...) \/			%
%             (?C1 B C2. (C = if B C1 C2) /\ B s1 ...) \/		%
%             (?C2 B C1. (C = if B C1 C2) /\ ~B s1 ...) \/		%
%             (?B C'. (C = while B C') /\ ~B s1 ... ) \/		%
%             (?C' B s2'. (C = while B C') /\ B s1 ...)			%
%  									%
% If C is specialized to some particular syntactic form, for example    %
% to "C1;C2", then most of the disjuncts in the conclusion become 	%
% just false because of the syntactic inequality of different commands. %
% These false can be simplified away by rewriting with the theorem	%
% distinct.  The disjuncts that match the command to which C is 	%
% specialized can also be simplified by rewriting with const11.  This   %
% changes equalities between similar commands, for example:		%
%									%
%    (C1 ; C2) = (C1' ; C2')						%
%									%
% to equalities between their coresponding constitutents:		%
%									%
%    C1 = C1'  /\  C2 = C2'						%
%									%
% These can then generally be used for substitution.			%
% --------------------------------------------------------------------- %

let SIMPLIFY = REWRITE_RULE (distinct . const11);;

% --------------------------------------------------------------------- %
% CASE_TAC : this is applicable to goals of the form:			%
%									%
%      TRANS C s1 s2 ==> P						%
%									%
% When applied to such a goal, the antecedant is matched to the general %
% case analysis theorem and a corresponding instance of its conclusion  %
% is obtained.  This is simplified using the SIMPLIFY rule described    %
% above and the result is assumed in stripped form.  Given this tactic, %
% the sequence of theorems given below are simple to prove.  The proofs %
% are fairly uniform; with a careful analysis of the regularities, one  %
% should be able to devise an automatic proof procedure for deriving    %
% sets of theorems of this type.					%
% --------------------------------------------------------------------- %

let CASE_TAC = DISCH_THEN 
               (STRIP_ASSUME_TAC o SIMPLIFY o ONCE_REWRITE_RULE[ecases]);;

% --------------------------------------------------------------------- %
% SKIP_THM : EVAL skip s1 s2 is provable only by the skip rule, which   %
% requires that s1 and s2 be the same state.				%
% --------------------------------------------------------------------- %

let SKIP_THM =
    prove_thm
    (`SKIP_THM`,
     "!s1 s2. EVAL skip s1 s2 = (s1 = s2)",
     REPEAT GEN_TAC THEN EQ_TAC THENL
     [CASE_TAC THEN ASM_REWRITE_TAC [];
      DISCH_THEN SUBST1_TAC THEN MAP_FIRST RULE_TAC rules]);;


% --------------------------------------------------------------------- %
% ASSIGN_THM : EVAL (V := E) s1 s2 is provable only by the assignment	%
% rule, which requires that s2 be the state s1 with V updated to E.	%
% --------------------------------------------------------------------- %

let ASSIGN_THM =
    prove_thm
    (`ASSIGN_THM`,
     "!s1 s2 V E. EVAL (V := E) s1 s2 = ((\v. ((v=V) => E s1 | s1 v)) = s2)",
     REPEAT GEN_TAC THEN EQ_TAC THENL
     [CASE_TAC THEN ASM_REWRITE_TAC [];
      DISCH_THEN (SUBST1_TAC o SYM) THEN MAP_FIRST RULE_TAC rules]);;


% --------------------------------------------------------------------- %
% SEQ_THM : EVAL (C1;C2) s1 s2 is provable only by the sequencing rule, %
% which requires that some intermediate state s3 exists such that C1    %
% in state s1 terminates in s3 and C3 in s3 terminates in s2.		%
% --------------------------------------------------------------------- %

let SEQ_THM =
    prove_thm
    (`SEQ_THM`,
     "!s1 s2 C1 C2.EVAL (C1;C2) s1 s2 = (?s3. EVAL C1 s1 s3 /\ EVAL C2 s3 s2)",
     REPEAT GEN_TAC THEN EQ_TAC THENL
     [CASE_TAC THEN EXISTS_TAC "s2':state" THEN ASM_REWRITE_TAC [];
      DISCH_THEN \th. MAP_FIRST RULE_TAC rules THEN MATCH_ACCEPT_TAC th]);;


% --------------------------------------------------------------------- %
% IF_T_THM : if B(s1) is true, then EVAL (if B C2 C2) s1 s2 is provable %
% only by the first conditional rule, which requires that C1 when	%
% evaluated in s1 terminates in s2.					%
% --------------------------------------------------------------------- %

let IF_T_THM =
    prove_thm
    (`IF_T_THM`,
     "!s1 s2 B C1 C2. B s1 ==> (EVAL (if B C1 C2) s1 s2 = EVAL C1 s1 s2)",
     REPEAT STRIP_TAC THEN EQ_TAC THENL
     [CASE_TAC THEN EVERY_ASSUM (TRY o SUBST_ALL_TAC) THENL
      [FIRST_ASSUM ACCEPT_TAC; RES_TAC];
      DISCH_TAC THEN MAP_FIRST RULE_TAC rules THEN FIRST_ASSUM ACCEPT_TAC]);;


% --------------------------------------------------------------------- %
% IF_F_THM : if B(s1) is false, then EVAL (if B C1 C2) s1 s2 is 	%
% provable only by the second conditional rule, which requires that C2	%
% when evaluated in s1 terminates in s2.				%
% --------------------------------------------------------------------- %

let IF_F_THM =
    prove_thm
    (`IF_F_THM`,
     "!s1 s2 B C1 C2. ~B s1 ==> (EVAL (if B C1 C2) s1 s2 = EVAL C2 s1 s2)",
     REPEAT STRIP_TAC THEN EQ_TAC THENL
     [CASE_TAC THEN EVERY_ASSUM (TRY o SUBST_ALL_TAC) THENL
      [RES_TAC; FIRST_ASSUM ACCEPT_TAC];
      DISCH_TAC THEN MAP_FIRST RULE_TAC (rev rules) THEN
      FIRST_ASSUM ACCEPT_TAC]);;


% --------------------------------------------------------------------- %
% WHILE_T_THM : if B(s1) is true, then EVAL (while B C) s1 s2 is 	%
% provable only by the corresponding while rule, which requires that 	%
% there is an intermediate state s3 such that C in state s1 terminates  %
% in s3, and while B do C in state s3 terminates in s2.			%
% --------------------------------------------------------------------- %

let WHILE_T_THM =
    prove_thm
    (`WHILE_T_THM`,
     "!s1 s2 B C.
      B s1 ==> (EVAL (while B C) s1 s2 =
                (?s3. EVAL C s1 s3 /\ EVAL (while B C) s3 s2))",
     REPEAT STRIP_TAC THEN EQ_TAC THENL
     [CASE_TAC THEN EVERY_ASSUM (TRY o SUBST_ALL_TAC) THENL
      [RES_TAC;
       EXISTS_TAC "s2':state" THEN CONJ_TAC THEN FIRST_ASSUM ACCEPT_TAC];
      STRIP_TAC THEN MAP_FIRST RULE_TAC (rev rules) THEN
      EXISTS_TAC "s3:state" THEN
      REPEAT CONJ_TAC THEN FIRST_ASSUM ACCEPT_TAC]);;


% --------------------------------------------------------------------- %
% WHILE_F_THM : if B(s1) is false, then EVAL (while B C) s1 s2 is	%
% provable only by the corresponding while rule, which requires that 	%
% s2 equals the original state s1					%
% --------------------------------------------------------------------- %

let WHILE_F_THM =
    prove_thm
    (`WHILE_F_THM`,
     "!s1 s2 B C. ~B s1 ==> (EVAL (while B C) s1 s2 = (s1 = s2))",
     REPEAT STRIP_TAC THEN EQ_TAC THENL
     [CASE_TAC THENL
      [CONV_TAC SYM_CONV THEN FIRST_ASSUM ACCEPT_TAC;
       EVERY_ASSUM (TRY o SUBST_ALL_TAC) THEN RES_TAC];
      DISCH_THEN (SUBST1_TAC o SYM) THEN MAP_FIRST RULE_TAC rules THEN
      FIRST_ASSUM ACCEPT_TAC]);;


% ===================================================================== %
% THEOREM: the operational semantics is deterministic.			%
%									%
% Given the suite of theorems proved above, this proof is relatively    %
% strightforward.  The standard proof is by structural induction on C,  %
% but the proof by rule induction given below gives rise to a slightly  %
% easier analysis in each case of the induction.  There are, however,   %
% more cases---one per rule, rather than one per constructor.		%
% ===================================================================== %

let DETERMINISTIC =
    prove_thm
    (`DETERMINISTIC`,
     "!C st1 st2. EVAL C st1 st2 ==> !st3. EVAL C st1 st3 ==> (st2 = st3)",
     RULE_INDUCT_TAC THEN REPEAT GEN_TAC THENL
     [REWRITE_TAC [SKIP_THM];
      REWRITE_TAC [ASSIGN_THM];
      PURE_ONCE_REWRITE_TAC [SEQ_THM] THEN STRIP_TAC THEN
      FIRST_ASSUM MATCH_MP_TAC THEN RES_TAC THEN ASM_REWRITE_TAC [];
      IMP_RES_TAC IF_T_THM THEN ASM_REWRITE_TAC [];
      IMP_RES_TAC IF_F_THM THEN ASM_REWRITE_TAC [];
      IMP_RES_TAC WHILE_F_THM THEN ASM_REWRITE_TAC [];
      IMP_RES_THEN (\th. PURE_ONCE_REWRITE_TAC [th]) WHILE_T_THM THEN
      STRIP_TAC THEN FIRST_ASSUM MATCH_MP_TAC THEN
      RES_TAC THEN ASM_REWRITE_TAC []]);;


% ===================================================================== %
% Definition of partial correctness and derivation of proof rules.	%
% ===================================================================== %

let SPEC_DEF =
    new_definition
    (`SPEC_DEF`,
     "SPEC P C Q = !s1 s2. (P s1 /\ EVAL C s1 s2) ==> Q s2");;

% --------------------------------------------------------------------- %
% Proof of the while rule in Hoare logic.				%
% --------------------------------------------------------------------- %

% --------------------------------------------------------------------- %
% In the following proofs, theorems of the form:			%
%									%
%     |- !y1...yn. (C x1 ... xn = C y1 ... yn) ==> tm[y1,...,yn]	%
%									%
% frequently arise, where C is one of the constructors of the data type %
% of commands.  The following theorem-tactic simplifies such theorems   %
% by specializing yi to xi and then removing the resulting trivially    %
% true antecedent.  The result is:					%
%									%
%     |- tm[x1,...,xn/y1,...,yn]					%
%									%
% which is passed to the theorem continuation function. The tactic just	%
% discards theorems not of the form shown above.   For the while proof  %
% given below, this has the effect of thinning out useless induction	%
% hypotheses of the form:						%
%									%
%     |- !B' C'. (C = while B' C') ==> tm[B',C']			%
% 									%
% These are just discarded.						%
% --------------------------------------------------------------------- %

let REFL_MP_THEN ttac th =
    (let tm = lhs(fst(dest_imp(snd(strip_forall(concl th))))) in 
     ttac (MATCH_MP th (REFL tm))) ? ALL_TAC;;

% --------------------------------------------------------------------- %
% The following lemma states that the condition B in while B C must be  %
% false upon termination of a while loop.  The proof is by a rule 	%
% induction specialized to the while rule cases.  We show that the set	%
% 									%
%    {(while B C,s1,s2) | ~(B s2)} U {(C,s1,s2) | ~(C = while B' C')}   %
%									%
% is closed under the rules for the evaluation relation. Note that this %
% formulation illustrates a general way of proving a property of some   %
% specific class of commands by rule induction.  One takes the union of %
% the set containing triples with the desired property and the set of   %
% all other triples whose command component is NOT an element of the    %
% class of commands of interest.					%
%									%
% The proof is trivial for all but the two while rules, since this set	%
% contains all triples (C,s1,s2) for which C is not a while command.  	%
% The subgoals corresponding to these cases are vacuously true, since 	%
% they are implications with antecedents of the form (C = while B' C'), %
% where C is a command syntactically distinct from any while command.	%
%									%
% Showing that the above set is closed under the two while rules is     %
% likewise trivial.  For the while axiom, we get ~(B s2) immediately 	%
% from the side condition. For the other while rule, the statement to	%
% prove is just one of the induction hypotheses; since RULE_INDUCT_TAC  %
% uses STRIP_ASSUME_TAC on this hypothesis, this subgoal is solved	%
% immediately.								%
% --------------------------------------------------------------------- %


let WHILE_LEMMA1 =
    TAC_PROOF
    (([], "!C s1 s2. EVAL C s1 s2 ==> !B' C'. (C = while B' C') ==> ~(B' s2)"),
     RULE_INDUCT_TAC THEN REWRITE_TAC (distinct . const11) THEN
     REPEAT GEN_TAC THEN DISCH_THEN (STRIP_THM_THEN SUBST_ALL_TAC) THEN
     FIRST_ASSUM ACCEPT_TAC);;

% --------------------------------------------------------------------- %
% The second lemma deals with the invariant part of the Hoare proof 	%
% rule for while commands.  We show that if P is an invariant of C, 	%
% then it is also an invariant of while B C.  The proof is essentially  %
% an induction on the number of applications of the evaluation rule for %
% while commands.  This is expressed as a rule induction, which 	%
% establishes that the set:						%
%									%
%    {(while B C,s1,s2) | P invariant of C ==> (P s1 ==> P s2)}		%
%									%
% is closed under the transition rules.  As in lemma 1, the rules for   %
% other kinds of commands are dealt with by taking the union of this 	%
% set with 								%
% 									%
%    {(C,s1,s2) | ~(C = while B' C')}					%
%									%
% Closure under evaluation rules other than the two rules for while is  %
% therefore trivial, as outlined in the comments to lemma 1 above. 	%
%									%
% The proof in fact proceeds by strong rule induction.  With ordinary   %
% rule induction, one obtains hypotheses that are too weak to imply the %
% desired conclusion in the step case of the while rule.  To see why, 	%
% try replacing strong by weak induction in the tactic proof below.	%
%									%
% Note that REFL_MP_THEN is used to simplify the induction hypotheses   %
% before adding them to the assumption list.  This avoids having the 	%
% assumptions in an awkward form (try using ASSUME_TAC instead). Note	%
% also that in the case of the while axiom, the states s1 and s2 are	%
% identical, so the corresponding subgoal is trivial and is solved by	%
% the rewriting step.							%
% --------------------------------------------------------------------- %

let WHILE_LEMMA2 =
    TAC_PROOF
    (([], "!C s1 s2. EVAL C s1 s2 
             ==>
           !B' C'. (C = while B' C') ==>
                   (!s1 s2. P s1 /\ B' s1 /\ EVAL C' s1 s2 ==> P s2) ==>
                   (P s1 ==> P s2)"),
     RULE_INDUCT_THEN sind (REFL_MP_THEN ASSUME_TAC) ASSUME_TAC THEN 
     REWRITE_TAC (distinct . const11) THEN REPEAT GEN_TAC THEN
     DISCH_THEN (STRIP_THM_THEN SUBST_ALL_TAC) THEN
     REPEAT STRIP_TAC THEN RES_TAC);;

% --------------------------------------------------------------------- %
% The proof rule for while commands in Hoare logic is:			%
%									%
%         |- {P /\ B} C {P}						%
%      ----------------------						%
%        |- {P} C {P /\ ~B}						%
% 									%
% Given the two lemmas proved above, it is trivial to prove this rule.  %
% The antecedent of the rule is just the assumption of invariance of P  %
% for C which occurs in lemma 2.  Note that REFL_MP_THEN is used to     %
% simplify the conclusions of the lemmas after one resolution step.	%
% --------------------------------------------------------------------- %

let WHILE =
    prove_thm
    (`WHILE`,
     "!P C. SPEC (\s. P s /\ (B s)) C P ==>
            SPEC P (while B C) (\s. P s /\ ~B s)",
     PURE_ONCE_REWRITE_TAC [SPEC_DEF] THEN
     CONV_TAC (ONCE_DEPTH_CONV BETA_CONV) THEN
     PURE_ONCE_REWRITE_TAC [SYM(SPEC_ALL CONJ_ASSOC)] THEN
     REPEAT STRIP_TAC THENL
     [IMP_RES_THEN (REFL_MP_THEN IMP_RES_TAC) WHILE_LEMMA2;
      IMP_RES_THEN (REFL_MP_THEN IMP_RES_TAC) WHILE_LEMMA1]);;

% --------------------------------------------------------------------- %
% End of example.							%
% --------------------------------------------------------------------- %

close_theory();;
quit();;