%

\begin{isabellebody}%

\def\isabellecontext{CTL}%

%

\isamarkupsubsection{Computation Tree Logic---CTL%

}

%

\begin{isamarkuptext}%

\label{sec:CTL}

The semantics of PDL only needs reflexive transitive closure.

Let us be adventurous and introduce a more expressive temporal operator.

We extend the datatype

\isa{formula} by a new constructor%

\end{isamarkuptext}%

\ \ \ \ \ \ \ \ \ \ \ \ \ \ \ \ \ \ {\isacharbar}\ AF\ formula%

\begin{isamarkuptext}%

\noindent

which stands for "always in the future":

on all paths, at some point the formula holds. Formalizing the notion of an infinite path is easy

in HOL: it is simply a function from \isa{nat} to \isa{state}.%

\end{isamarkuptext}%

\isacommand{constdefs}\ Paths\ {\isacharcolon}{\isacharcolon}\ {\isachardoublequote}state\ {\isasymRightarrow}\ {\isacharparenleft}nat\ {\isasymRightarrow}\ state{\isacharparenright}set{\isachardoublequote}\isanewline

\ \ \ \ \ \ \ \ \ {\isachardoublequote}Paths\ s\ {\isasymequiv}\ {\isacharbraceleft}p{\isachardot}\ s\ {\isacharequal}\ p\ {\isadigit{0}}\ {\isasymand}\ {\isacharparenleft}{\isasymforall}i{\isachardot}\ {\isacharparenleft}p\ i{\isacharcomma}\ p{\isacharparenleft}i{\isacharplus}{\isadigit{1}}{\isacharparenright}{\isacharparenright}\ {\isasymin}\ M{\isacharparenright}{\isacharbraceright}{\isachardoublequote}%

\begin{isamarkuptext}%

\noindent

This definition allows a very succinct statement of the semantics of \isa{AF}:

\footnote{Do not be misled: neither datatypes nor recursive functions can be

extended by new constructors or equations. This is just a trick of the

presentation. In reality one has to define a new datatype and a new function.}%

\end{isamarkuptext}%

{\isachardoublequote}s\ {\isasymTurnstile}\ AF\ f\ \ \ \ {\isacharequal}\ {\isacharparenleft}{\isasymforall}p\ {\isasymin}\ Paths\ s{\isachardot}\ {\isasymexists}i{\isachardot}\ p\ i\ {\isasymTurnstile}\ f{\isacharparenright}{\isachardoublequote}%

\begin{isamarkuptext}%

\noindent

Model checking \isa{AF} involves a function which

is just complicated enough to warrant a separate definition:%

\end{isamarkuptext}%

\isacommand{constdefs}\ af\ {\isacharcolon}{\isacharcolon}\ {\isachardoublequote}state\ set\ {\isasymRightarrow}\ state\ set\ {\isasymRightarrow}\ state\ set{\isachardoublequote}\isanewline

\ \ \ \ \ \ \ \ \ {\isachardoublequote}af\ A\ T\ {\isasymequiv}\ A\ {\isasymunion}\ {\isacharbraceleft}s{\isachardot}\ {\isasymforall}t{\isachardot}\ {\isacharparenleft}s{\isacharcomma}\ t{\isacharparenright}\ {\isasymin}\ M\ {\isasymlongrightarrow}\ t\ {\isasymin}\ T{\isacharbraceright}{\isachardoublequote}%

\begin{isamarkuptext}%

\noindent

Now we define \isa{mc\ {\isacharparenleft}AF\ f{\isacharparenright}} as the least set \isa{T} that includes

\isa{mc\ f} and all states all of whose direct successors are in \isa{T}:%

\end{isamarkuptext}%

{\isachardoublequote}mc{\isacharparenleft}AF\ f{\isacharparenright}\ \ \ \ {\isacharequal}\ lfp{\isacharparenleft}af{\isacharparenleft}mc\ f{\isacharparenright}{\isacharparenright}{\isachardoublequote}%

\begin{isamarkuptext}%

\noindent

Because \isa{af} is monotone in its second argument (and also its first, but

that is irrelevant) \isa{af\ A} has a least fixed point:%

\end{isamarkuptext}%

\isacommand{lemma}\ mono{\isacharunderscore}af{\isacharcolon}\ {\isachardoublequote}mono{\isacharparenleft}af\ A{\isacharparenright}{\isachardoublequote}\isanewline

\isacommand{apply}{\isacharparenleft}simp\ add{\isacharcolon}\ mono{\isacharunderscore}def\ af{\isacharunderscore}def{\isacharparenright}\isanewline

\isacommand{apply}\ blast\isanewline

\isacommand{done}%

\begin{isamarkuptext}%

All we need to prove now is  \isa{mc\ {\isacharparenleft}AF\ f{\isacharparenright}\ {\isacharequal}\ {\isacharbraceleft}s{\isachardot}\ s\ {\isasymTurnstile}\ AF\ f{\isacharbraceright}}, which states

that \isa{mc} and \isa{{\isasymTurnstile}} agree for \isa{AF}\@.

This time we prove the two inclusions separately, starting

with the easy one:%

\end{isamarkuptext}%

\isacommand{theorem}\ AF{\isacharunderscore}lemma{\isadigit{1}}{\isacharcolon}\isanewline

\ \ {\isachardoublequote}lfp{\isacharparenleft}af\ A{\isacharparenright}\ {\isasymsubseteq}\ {\isacharbraceleft}s{\isachardot}\ {\isasymforall}\ p\ {\isasymin}\ Paths\ s{\isachardot}\ {\isasymexists}\ i{\isachardot}\ p\ i\ {\isasymin}\ A{\isacharbraceright}{\isachardoublequote}%

\begin{isamarkuptxt}%

\noindent

In contrast to the analogous property for \isa{EF}, and just

for a change, we do not use fixed point induction but a weaker theorem,

\isa{lfp{\isacharunderscore}lowerbound}:

\begin{isabelle}%

\ \ \ \ \ f\ S\ {\isasymsubseteq}\ S\ {\isasymLongrightarrow}\ lfp\ f\ {\isasymsubseteq}\ S%

\end{isabelle}

The instance of the premise \isa{f\ S\ {\isasymsubseteq}\ S} is proved pointwise,

a decision that clarification takes for us:%

\end{isamarkuptxt}%

\isacommand{apply}{\isacharparenleft}rule\ lfp{\isacharunderscore}lowerbound{\isacharparenright}\isanewline

\isacommand{apply}{\isacharparenleft}clarsimp\ simp\ add{\isacharcolon}\ af{\isacharunderscore}def\ Paths{\isacharunderscore}def{\isacharparenright}%

\begin{isamarkuptxt}%

\begin{isabelle}%

\ {\isadigit{1}}{\isachardot}\ {\isasymAnd}p{\isachardot}\ {\isasymlbrakk}p\ {\isadigit{0}}\ {\isasymin}\ A\ {\isasymor}\isanewline

\ \ \ \ \ \ \ \ \ {\isacharparenleft}{\isasymforall}t{\isachardot}\ {\isacharparenleft}p\ {\isadigit{0}}{\isacharcomma}\ t{\isacharparenright}\ {\isasymin}\ M\ {\isasymlongrightarrow}\isanewline

\ \ \ \ \ \ \ \ \ \ \ \ \ \ {\isacharparenleft}{\isasymforall}p{\isachardot}\ t\ {\isacharequal}\ p\ {\isadigit{0}}\ {\isasymand}\ {\isacharparenleft}{\isasymforall}i{\isachardot}\ {\isacharparenleft}p\ i{\isacharcomma}\ p\ {\isacharparenleft}Suc\ i{\isacharparenright}{\isacharparenright}\ {\isasymin}\ M{\isacharparenright}\ {\isasymlongrightarrow}\isanewline

\ \ \ \ \ \ \ \ \ \ \ \ \ \ \ \ \ \ \ {\isacharparenleft}{\isasymexists}i{\isachardot}\ p\ i\ {\isasymin}\ A{\isacharparenright}{\isacharparenright}{\isacharparenright}{\isacharsemicolon}\isanewline

\ \ \ \ \ \ \ \ \ \ \ {\isasymforall}i{\isachardot}\ {\isacharparenleft}p\ i{\isacharcomma}\ p\ {\isacharparenleft}Suc\ i{\isacharparenright}{\isacharparenright}\ {\isasymin}\ M{\isasymrbrakk}\isanewline

\ \ \ \ \ \ \ \ {\isasymLongrightarrow}\ {\isasymexists}i{\isachardot}\ p\ i\ {\isasymin}\ A%

\end{isabelle}

Now we eliminate the disjunction. The case \isa{p\ {\isadigit{0}}\ {\isasymin}\ A} is trivial:%

\end{isamarkuptxt}%

\isacommand{apply}{\isacharparenleft}erule\ disjE{\isacharparenright}\isanewline

\ \isacommand{apply}{\isacharparenleft}blast{\isacharparenright}%

\begin{isamarkuptxt}%

\noindent

In the other case we set \isa{t} to \isa{p\ {\isadigit{1}}} and simplify matters:%

\end{isamarkuptxt}%

\isacommand{apply}{\isacharparenleft}erule{\isacharunderscore}tac\ x\ {\isacharequal}\ {\isachardoublequote}p\ {\isadigit{1}}{\isachardoublequote}\ \isakeyword{in}\ allE{\isacharparenright}\isanewline

\isacommand{apply}{\isacharparenleft}clarsimp{\isacharparenright}%

\begin{isamarkuptxt}%

\begin{isabelle}%

\ {\isadigit{1}}{\isachardot}\ {\isasymAnd}p{\isachardot}\ {\isasymlbrakk}{\isasymforall}i{\isachardot}\ {\isacharparenleft}p\ i{\isacharcomma}\ p\ {\isacharparenleft}Suc\ i{\isacharparenright}{\isacharparenright}\ {\isasymin}\ M{\isacharsemicolon}\isanewline

\ \ \ \ \ \ \ \ \ \ \ {\isasymforall}pa{\isachardot}\ p\ {\isadigit{1}}\ {\isacharequal}\ pa\ {\isadigit{0}}\ {\isasymand}\ {\isacharparenleft}{\isasymforall}i{\isachardot}\ {\isacharparenleft}pa\ i{\isacharcomma}\ pa\ {\isacharparenleft}Suc\ i{\isacharparenright}{\isacharparenright}\ {\isasymin}\ M{\isacharparenright}\ {\isasymlongrightarrow}\isanewline

\ \ \ \ \ \ \ \ \ \ \ \ \ \ \ \ {\isacharparenleft}{\isasymexists}i{\isachardot}\ pa\ i\ {\isasymin}\ A{\isacharparenright}{\isasymrbrakk}\isanewline

\ \ \ \ \ \ \ \ {\isasymLongrightarrow}\ {\isasymexists}i{\isachardot}\ p\ i\ {\isasymin}\ A%

\end{isabelle}

It merely remains to set \isa{pa} to \isa{{\isasymlambda}i{\isachardot}\ p\ {\isacharparenleft}i\ {\isacharplus}\ {\isadigit{1}}{\isacharparenright}}, i.e.\ \isa{p} without its

first element. The rest is practically automatic:%

\end{isamarkuptxt}%

\isacommand{apply}{\isacharparenleft}erule{\isacharunderscore}tac\ x\ {\isacharequal}\ {\isachardoublequote}{\isasymlambda}i{\isachardot}\ p{\isacharparenleft}i{\isacharplus}{\isadigit{1}}{\isacharparenright}{\isachardoublequote}\ \isakeyword{in}\ allE{\isacharparenright}\isanewline

\isacommand{apply}\ simp\isanewline

\isacommand{apply}\ blast\isanewline

\isacommand{done}%

\begin{isamarkuptext}%

The opposite inclusion is proved by contradiction: if some state

\isa{s} is not in \isa{lfp\ {\isacharparenleft}af\ A{\isacharparenright}}, then we can construct an

infinite \isa{A}-avoiding path starting from \isa{s}. The reason is

that by unfolding \isa{lfp} we find that if \isa{s} is not in

\isa{lfp\ {\isacharparenleft}af\ A{\isacharparenright}}, then \isa{s} is not in \isa{A} and there is a

direct successor of \isa{s} that is again not in \isa{lfp\ {\isacharparenleft}af\ A{\isacharparenright}}. Iterating this argument yields the promised infinite

\isa{A}-avoiding path. Let us formalize this sketch.

The one-step argument in the sketch above

is proved by a variant of contraposition:%

\end{isamarkuptext}%

\isacommand{lemma}\ not{\isacharunderscore}in{\isacharunderscore}lfp{\isacharunderscore}afD{\isacharcolon}\isanewline

\ {\isachardoublequote}s\ {\isasymnotin}\ lfp{\isacharparenleft}af\ A{\isacharparenright}\ {\isasymLongrightarrow}\ s\ {\isasymnotin}\ A\ {\isasymand}\ {\isacharparenleft}{\isasymexists}\ t{\isachardot}\ {\isacharparenleft}s{\isacharcomma}t{\isacharparenright}{\isasymin}M\ {\isasymand}\ t\ {\isasymnotin}\ lfp{\isacharparenleft}af\ A{\isacharparenright}{\isacharparenright}{\isachardoublequote}\isanewline

\isacommand{apply}{\isacharparenleft}erule\ contrapos{\isacharunderscore}np{\isacharparenright}\isanewline

\isacommand{apply}{\isacharparenleft}rule\ ssubst{\isacharbrackleft}OF\ lfp{\isacharunderscore}unfold{\isacharbrackleft}OF\ mono{\isacharunderscore}af{\isacharbrackright}{\isacharbrackright}{\isacharparenright}\isanewline

\isacommand{apply}{\isacharparenleft}simp\ add{\isacharcolon}af{\isacharunderscore}def{\isacharparenright}\isanewline

\isacommand{done}%

\begin{isamarkuptext}%

\noindent

We assume the negation of the conclusion and prove \isa{s\ {\isasymin}\ lfp\ {\isacharparenleft}af\ A{\isacharparenright}}.

Unfolding \isa{lfp} once and

simplifying with the definition of \isa{af} finishes the proof.

Now we iterate this process. The following construction of the desired

path is parameterized by a predicate \isa{P} that should hold along the path:%

\end{isamarkuptext}%

\isacommand{consts}\ path\ {\isacharcolon}{\isacharcolon}\ {\isachardoublequote}state\ {\isasymRightarrow}\ {\isacharparenleft}state\ {\isasymRightarrow}\ bool{\isacharparenright}\ {\isasymRightarrow}\ {\isacharparenleft}nat\ {\isasymRightarrow}\ state{\isacharparenright}{\isachardoublequote}\isanewline

\isacommand{primrec}\isanewline

{\isachardoublequote}path\ s\ P\ {\isadigit{0}}\ {\isacharequal}\ s{\isachardoublequote}\isanewline

{\isachardoublequote}path\ s\ P\ {\isacharparenleft}Suc\ n{\isacharparenright}\ {\isacharequal}\ {\isacharparenleft}SOME\ t{\isachardot}\ {\isacharparenleft}path\ s\ P\ n{\isacharcomma}t{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P\ t{\isacharparenright}{\isachardoublequote}%

\begin{isamarkuptext}%

\noindent

Element \isa{n\ {\isacharplus}\ {\isadigit{1}}} on this path is some arbitrary successor

\isa{t} of element \isa{n} such that \isa{P\ t} holds.  Remember that \isa{SOME\ t{\isachardot}\ R\ t}

is some arbitrary but fixed \isa{t} such that \isa{R\ t} holds (see \S\ref{sec:SOME}). Of

course, such a \isa{t} need not exist, but that is of no

concern to us since we will only use \isa{path} when a

suitable \isa{t} does exist.

Let us show that if each state \isa{s} that satisfies \isa{P}

has a successor that again satisfies \isa{P}, then there exists an infinite \isa{P}-path:%

\end{isamarkuptext}%

\isacommand{lemma}\ infinity{\isacharunderscore}lemma{\isacharcolon}\isanewline

\ \ {\isachardoublequote}{\isasymlbrakk}\ P\ s{\isacharsemicolon}\ {\isasymforall}s{\isachardot}\ P\ s\ {\isasymlongrightarrow}\ {\isacharparenleft}{\isasymexists}\ t{\isachardot}\ {\isacharparenleft}s{\isacharcomma}t{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P\ t{\isacharparenright}\ {\isasymrbrakk}\ {\isasymLongrightarrow}\isanewline

\ \ \ {\isasymexists}p{\isasymin}Paths\ s{\isachardot}\ {\isasymforall}i{\isachardot}\ P{\isacharparenleft}p\ i{\isacharparenright}{\isachardoublequote}%

\begin{isamarkuptxt}%

\noindent

First we rephrase the conclusion slightly because we need to prove both the path property

and the fact that \isa{P} holds simultaneously:%

\end{isamarkuptxt}%

\isacommand{apply}{\isacharparenleft}subgoal{\isacharunderscore}tac\ {\isachardoublequote}{\isasymexists}p{\isachardot}\ s\ {\isacharequal}\ p\ {\isadigit{0}}\ {\isasymand}\ {\isacharparenleft}{\isasymforall}i{\isachardot}\ {\isacharparenleft}p\ i{\isacharcomma}p{\isacharparenleft}i{\isacharplus}{\isadigit{1}}{\isacharparenright}{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P{\isacharparenleft}p\ i{\isacharparenright}{\isacharparenright}{\isachardoublequote}{\isacharparenright}%

\begin{isamarkuptxt}%

\noindent

From this proposition the original goal follows easily:%

\end{isamarkuptxt}%

\ \isacommand{apply}{\isacharparenleft}simp\ add{\isacharcolon}Paths{\isacharunderscore}def{\isacharcomma}\ blast{\isacharparenright}%

\begin{isamarkuptxt}%

\noindent

The new subgoal is proved by providing the witness \isa{path\ s\ P} for \isa{p}:%

\end{isamarkuptxt}%

\isacommand{apply}{\isacharparenleft}rule{\isacharunderscore}tac\ x\ {\isacharequal}\ {\isachardoublequote}path\ s\ P{\isachardoublequote}\ \isakeyword{in}\ exI{\isacharparenright}\isanewline

\isacommand{apply}{\isacharparenleft}clarsimp{\isacharparenright}%

\begin{isamarkuptxt}%

\noindent

After simplification and clarification the subgoal has the following compact form

\begin{isabelle}%

\ {\isadigit{1}}{\isachardot}\ {\isasymAnd}i{\isachardot}\ {\isasymlbrakk}P\ s{\isacharsemicolon}\ {\isasymforall}s{\isachardot}\ P\ s\ {\isasymlongrightarrow}\ {\isacharparenleft}{\isasymexists}t{\isachardot}\ {\isacharparenleft}s{\isacharcomma}\ t{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P\ t{\isacharparenright}{\isasymrbrakk}\isanewline

\ \ \ \ \ \ \ \ {\isasymLongrightarrow}\ {\isacharparenleft}path\ s\ P\ i{\isacharcomma}\ SOME\ t{\isachardot}\ {\isacharparenleft}path\ s\ P\ i{\isacharcomma}\ t{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P\ t{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\isanewline

\ \ \ \ \ \ \ \ \ \ \ P\ {\isacharparenleft}path\ s\ P\ i{\isacharparenright}%

\end{isabelle}

and invites a proof by induction on \isa{i}:%

\end{isamarkuptxt}%

\isacommand{apply}{\isacharparenleft}induct{\isacharunderscore}tac\ i{\isacharparenright}\isanewline

\ \isacommand{apply}{\isacharparenleft}simp{\isacharparenright}%

\begin{isamarkuptxt}%

\noindent

After simplification the base case boils down to

\begin{isabelle}%

\ {\isadigit{1}}{\isachardot}\ {\isasymlbrakk}P\ s{\isacharsemicolon}\ {\isasymforall}s{\isachardot}\ P\ s\ {\isasymlongrightarrow}\ {\isacharparenleft}{\isasymexists}t{\isachardot}\ {\isacharparenleft}s{\isacharcomma}\ t{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P\ t{\isacharparenright}{\isasymrbrakk}\isanewline

\ \ \ \ {\isasymLongrightarrow}\ {\isacharparenleft}s{\isacharcomma}\ SOME\ t{\isachardot}\ {\isacharparenleft}s{\isacharcomma}\ t{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P\ t{\isacharparenright}\ {\isasymin}\ M%

\end{isabelle}

The conclusion looks exceedingly trivial: after all, \isa{t} is chosen such that \isa{{\isacharparenleft}s{\isacharcomma}\ t{\isacharparenright}\ {\isasymin}\ M}

holds. However, we first have to show that such a \isa{t} actually exists! This reasoning

is embodied in the theorem \isa{someI{\isadigit{2}}{\isacharunderscore}ex}:

\begin{isabelle}%

\ \ \ \ \ {\isasymlbrakk}{\isasymexists}a{\isachardot}\ {\isacharquery}P\ a{\isacharsemicolon}\ {\isasymAnd}x{\isachardot}\ {\isacharquery}P\ x\ {\isasymLongrightarrow}\ {\isacharquery}Q\ x{\isasymrbrakk}\ {\isasymLongrightarrow}\ {\isacharquery}Q\ {\isacharparenleft}SOME\ x{\isachardot}\ {\isacharquery}P\ x{\isacharparenright}%

\end{isabelle}

When we apply this theorem as an introduction rule, \isa{{\isacharquery}P\ x} becomes

\isa{{\isacharparenleft}s{\isacharcomma}\ x{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P\ x} and \isa{{\isacharquery}Q\ x} becomes \isa{{\isacharparenleft}s{\isacharcomma}\ x{\isacharparenright}\ {\isasymin}\ M} and we have to prove

two subgoals: \isa{{\isasymexists}a{\isachardot}\ {\isacharparenleft}s{\isacharcomma}\ a{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P\ a}, which follows from the assumptions, and

\isa{{\isacharparenleft}s{\isacharcomma}\ x{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P\ x\ {\isasymLongrightarrow}\ {\isacharparenleft}s{\isacharcomma}\ x{\isacharparenright}\ {\isasymin}\ M}, which is trivial. Thus it is not surprising that

\isa{fast} can prove the base case quickly:%

\end{isamarkuptxt}%

\ \isacommand{apply}{\isacharparenleft}fast\ intro{\isacharcolon}someI{\isadigit{2}}{\isacharunderscore}ex{\isacharparenright}%

\begin{isamarkuptxt}%

\noindent

What is worth noting here is that we have used \isa{fast} rather than

\isa{blast}.  The reason is that \isa{blast} would fail because it cannot

cope with \isa{someI{\isadigit{2}}{\isacharunderscore}ex}: unifying its conclusion with the current

subgoal is nontrivial because of the nested schematic variables. For

efficiency reasons \isa{blast} does not even attempt such unifications.

Although \isa{fast} can in principle cope with complicated unification

problems, in practice the number of unifiers arising is often prohibitive and

the offending rule may need to be applied explicitly rather than

automatically. This is what happens in the step case.

The induction step is similar, but more involved, because now we face nested

occurrences of \isa{SOME}. As a result, \isa{fast} is no longer able to

solve the subgoal and we apply \isa{someI{\isadigit{2}}{\isacharunderscore}ex} by hand.  We merely

show the proof commands but do not describe the details:%

\end{isamarkuptxt}%

\isacommand{apply}{\isacharparenleft}simp{\isacharparenright}\isanewline

\isacommand{apply}{\isacharparenleft}rule\ someI{\isadigit{2}}{\isacharunderscore}ex{\isacharparenright}\isanewline

\ \isacommand{apply}{\isacharparenleft}blast{\isacharparenright}\isanewline

\isacommand{apply}{\isacharparenleft}rule\ someI{\isadigit{2}}{\isacharunderscore}ex{\isacharparenright}\isanewline

\ \isacommand{apply}{\isacharparenleft}blast{\isacharparenright}\isanewline

\isacommand{apply}{\isacharparenleft}blast{\isacharparenright}\isanewline

\isacommand{done}%

\begin{isamarkuptext}%

Function \isa{path} has fulfilled its purpose now and can be forgotten.

It was merely defined to provide the witness in the proof of the

\isa{infinity{\isacharunderscore}lemma}. Aficionados of minimal proofs might like to know

that we could have given the witness without having to define a new function:

the term

\begin{isabelle}%

\ \ \ \ \ nat{\isacharunderscore}rec\ s\ {\isacharparenleft}{\isasymlambda}n\ t{\isachardot}\ SOME\ u{\isachardot}\ {\isacharparenleft}t{\isacharcomma}\ u{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ P\ u{\isacharparenright}%

\end{isabelle}

is extensionally equal to \isa{path\ s\ P},

where \isa{nat{\isacharunderscore}rec} is the predefined primitive recursor on \isa{nat}.%

\end{isamarkuptext}%

%

\begin{isamarkuptext}%

At last we can prove the opposite direction of \isa{AF{\isacharunderscore}lemma{\isadigit{1}}}:%

\end{isamarkuptext}%

\isacommand{theorem}\ AF{\isacharunderscore}lemma{\isadigit{2}}{\isacharcolon}\ {\isachardoublequote}{\isacharbraceleft}s{\isachardot}\ {\isasymforall}\ p\ {\isasymin}\ Paths\ s{\isachardot}\ {\isasymexists}\ i{\isachardot}\ p\ i\ {\isasymin}\ A{\isacharbraceright}\ {\isasymsubseteq}\ lfp{\isacharparenleft}af\ A{\isacharparenright}{\isachardoublequote}%

\begin{isamarkuptxt}%

\noindent

The proof is again pointwise and then by contraposition:%

\end{isamarkuptxt}%

\isacommand{apply}{\isacharparenleft}rule\ subsetI{\isacharparenright}\isanewline

\isacommand{apply}{\isacharparenleft}erule\ contrapos{\isacharunderscore}pp{\isacharparenright}\isanewline

\isacommand{apply}\ simp%

\begin{isamarkuptxt}%

\begin{isabelle}%

\ {\isadigit{1}}{\isachardot}\ {\isasymAnd}x{\isachardot}\ x\ {\isasymnotin}\ lfp\ {\isacharparenleft}af\ A{\isacharparenright}\ {\isasymLongrightarrow}\ {\isasymexists}p{\isasymin}Paths\ x{\isachardot}\ {\isasymforall}i{\isachardot}\ p\ i\ {\isasymnotin}\ A%

\end{isabelle}

Applying the \isa{infinity{\isacharunderscore}lemma} as a destruction rule leaves two subgoals, the second

premise of \isa{infinity{\isacharunderscore}lemma} and the original subgoal:%

\end{isamarkuptxt}%

\isacommand{apply}{\isacharparenleft}drule\ infinity{\isacharunderscore}lemma{\isacharparenright}%

\begin{isamarkuptxt}%

\begin{isabelle}%

\ {\isadigit{1}}{\isachardot}\ {\isasymAnd}x{\isachardot}\ {\isasymforall}s{\isachardot}\ s\ {\isasymnotin}\ lfp\ {\isacharparenleft}af\ A{\isacharparenright}\ {\isasymlongrightarrow}\ {\isacharparenleft}{\isasymexists}t{\isachardot}\ {\isacharparenleft}s{\isacharcomma}\ t{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ t\ {\isasymnotin}\ lfp\ {\isacharparenleft}af\ A{\isacharparenright}{\isacharparenright}\isanewline

\ {\isadigit{2}}{\isachardot}\ {\isasymAnd}x{\isachardot}\ {\isasymexists}p{\isasymin}Paths\ x{\isachardot}\ {\isasymforall}i{\isachardot}\ p\ i\ {\isasymnotin}\ lfp\ {\isacharparenleft}af\ A{\isacharparenright}\ {\isasymLongrightarrow}\isanewline

\ \ \ \ \ \ \ \ {\isasymexists}p{\isasymin}Paths\ x{\isachardot}\ {\isasymforall}i{\isachardot}\ p\ i\ {\isasymnotin}\ A%

\end{isabelle}

Both are solved automatically:%

\end{isamarkuptxt}%

\ \isacommand{apply}{\isacharparenleft}auto\ dest{\isacharcolon}not{\isacharunderscore}in{\isacharunderscore}lfp{\isacharunderscore}afD{\isacharparenright}\isanewline

\isacommand{done}%

\begin{isamarkuptext}%

If you find these proofs too complicated, we recommend that you read

\S\ref{sec:CTL-revisited}, where we show how inductive definitions lead to

simpler arguments.

The main theorem is proved as for PDL, except that we also derive the

necessary equality \isa{lfp{\isacharparenleft}af\ A{\isacharparenright}\ {\isacharequal}\ {\isachardot}{\isachardot}{\isachardot}} by combining

\isa{AF{\isacharunderscore}lemma{\isadigit{1}}} and \isa{AF{\isacharunderscore}lemma{\isadigit{2}}} on the spot:%

\end{isamarkuptext}%

\isacommand{theorem}\ {\isachardoublequote}mc\ f\ {\isacharequal}\ {\isacharbraceleft}s{\isachardot}\ s\ {\isasymTurnstile}\ f{\isacharbraceright}{\isachardoublequote}\isanewline

\isacommand{apply}{\isacharparenleft}induct{\isacharunderscore}tac\ f{\isacharparenright}\isanewline

\isacommand{apply}{\isacharparenleft}auto\ simp\ add{\isacharcolon}\ EF{\isacharunderscore}lemma\ equalityI{\isacharbrackleft}OF\ AF{\isacharunderscore}lemma{\isadigit{1}}\ AF{\isacharunderscore}lemma{\isadigit{2}}{\isacharbrackright}{\isacharparenright}\isanewline

\isacommand{done}%

\begin{isamarkuptext}%

The language defined above is not quite CTL\@. The latter also includes an

until-operator \isa{EU\ f\ g} with semantics ``there exists a path

where \isa{f} is true until \isa{g} becomes true''. With the help

of an auxiliary function%

\end{isamarkuptext}%

\isacommand{consts}\ until{\isacharcolon}{\isacharcolon}\ {\isachardoublequote}state\ set\ {\isasymRightarrow}\ state\ set\ {\isasymRightarrow}\ state\ {\isasymRightarrow}\ state\ list\ {\isasymRightarrow}\ bool{\isachardoublequote}\isanewline

\isacommand{primrec}\isanewline

{\isachardoublequote}until\ A\ B\ s\ {\isacharbrackleft}{\isacharbrackright}\ \ \ \ {\isacharequal}\ {\isacharparenleft}s\ {\isasymin}\ B{\isacharparenright}{\isachardoublequote}\isanewline

{\isachardoublequote}until\ A\ B\ s\ {\isacharparenleft}t{\isacharhash}p{\isacharparenright}\ {\isacharequal}\ {\isacharparenleft}s\ {\isasymin}\ A\ {\isasymand}\ {\isacharparenleft}s{\isacharcomma}t{\isacharparenright}\ {\isasymin}\ M\ {\isasymand}\ until\ A\ B\ t\ p{\isacharparenright}{\isachardoublequote}%

\begin{isamarkuptext}%

\noindent

the semantics of \isa{EU} is straightforward:

\begin{isabelle}%

\ \ \ \ \ s\ {\isasymTurnstile}\ EU\ f\ g\ {\isacharequal}\ {\isacharparenleft}{\isasymexists}p{\isachardot}\ until\ A\ B\ s\ p{\isacharparenright}%

\end{isabelle}

Note that \isa{EU} is not definable in terms of the other operators!

Model checking \isa{EU} is again a least fixed point construction:

\begin{isabelle}%

\ \ \ \ \ mc{\isacharparenleft}EU\ f\ g{\isacharparenright}\ {\isacharequal}\ lfp{\isacharparenleft}{\isasymlambda}T{\isachardot}\ mc\ g\ {\isasymunion}\ mc\ f\ {\isasyminter}\ {\isacharparenleft}M{\isasyminverse}\ {\isacharbackquote}{\isacharbackquote}\ T{\isacharparenright}{\isacharparenright}%

\end{isabelle}

\begin{exercise}

Extend the datatype of formulae by the above until operator

and prove the equivalence between semantics and model checking, i.e.\ that

\begin{isabelle}%

\ \ \ \ \ mc\ {\isacharparenleft}EU\ f\ g{\isacharparenright}\ {\isacharequal}\ {\isacharbraceleft}s{\isachardot}\ s\ {\isasymTurnstile}\ EU\ f\ g{\isacharbraceright}%

\end{isabelle}

%For readability you may want to annotate {term EU} with its customary syntax

%{text[display]"| EU formula formula    E[_ U _]"}

%which enables you to read and write {text"E[f U g]"} instead of {term"EU f g"}.

\end{exercise}

For more CTL exercises see, for example, Huth and Ryan \cite{Huth-Ryan-book}.%

\end{isamarkuptext}%

%

\begin{isamarkuptext}%

Let us close this section with a few words about the executability of our model checkers.

It is clear that if all sets are finite, they can be represented as lists and the usual

set operations are easily implemented. Only \isa{lfp} requires a little thought.

Fortunately, the HOL Library%

\footnote{See theory \isa{While_Combinator_Example}.}

provides a theorem stating that 

in the case of finite sets and a monotone function~\isa{F},

the value of \isa{lfp\ F} can be computed by iterated application of \isa{F} to~\isa{{\isacharbraceleft}{\isacharbraceright}} until

a fixed point is reached. It is actually possible to generate executable functional programs

from HOL definitions, but that is beyond the scope of the tutorial.%

\end{isamarkuptext}%

\end{isabellebody}%

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