%% $Id$

\chapter{Theorems and Forward Proof}

\index{theorems|(}

Theorems, which represent the axioms, theorems and rules of

object-logics, have type \mltydx{thm}.  This chapter begins by

describing operations that print theorems and that join them in

forward proof.  Most theorem operations are intended for advanced

applications, such as programming new proof procedures.  Many of these

operations refer to signatures, certified terms and certified types,

which have the \ML{} types {\tt Sign.sg}, {\tt cterm} and {\tt ctyp}

and are discussed in Chapter~\ref{theories}.  Beginning users should

ignore such complexities --- and skip all but the first section of

this chapter.

The theorem operations do not print error messages.  Instead, they raise

exception~\xdx{THM}\@.  Use \ttindex{print_exn} to display

exceptions nicely:

\begin{ttbox} 

allI RS mp  handle e => print_exn e;

{\out Exception THM raised:}

{\out RSN: no unifiers -- premise 1}

{\out (!!x. ?P(x)) ==> ALL x. ?P(x)}

{\out [| ?P --> ?Q; ?P |] ==> ?Q}

{\out}

{\out uncaught exception THM}

\end{ttbox}

\section{Basic operations on theorems}

\subsection{Pretty-printing a theorem}

\index{theorems!printing of}

\begin{ttbox} 

prth          : thm -> thm

prths         : thm list -> thm list

prthq         : thm Sequence.seq -> thm Sequence.seq

print_thm     : thm -> unit

print_goals   : int -> thm -> unit

string_of_thm : thm -> string

\end{ttbox}

The first three commands are for interactive use.  They are identity

functions that display, then return, their argument.  The \ML{} identifier

{\tt it} will refer to the value just displayed.

The others are for use in programs.  Functions with result type {\tt unit}

are convenient for imperative programming.

\begin{ttdescription}

\item[\ttindexbold{prth} {\it thm}]  

prints {\it thm\/} at the terminal.

\item[\ttindexbold{prths} {\it thms}]  

prints {\it thms}, a list of theorems.

\item[\ttindexbold{prthq} {\it thmq}]  

prints {\it thmq}, a sequence of theorems.  It is useful for inspecting

the output of a tactic.

\item[\ttindexbold{print_thm} {\it thm}]  

prints {\it thm\/} at the terminal.

\item[\ttindexbold{print_goals} {\it limit\/} {\it thm}]  

prints {\it thm\/} in goal style, with the premises as subgoals.  It prints

at most {\it limit\/} subgoals.  The subgoal module calls {\tt print_goals}

to display proof states.

\item[\ttindexbold{string_of_thm} {\it thm}]  

converts {\it thm\/} to a string.

\end{ttdescription}

\subsection{Forward proof: joining rules by resolution}

\index{theorems!joining by resolution}

\index{resolution}\index{forward proof}

\begin{ttbox} 

RSN : thm * (int * thm) -> thm                 \hfill{\bf infix}

RS  : thm * thm -> thm                         \hfill{\bf infix}

MRS : thm list * thm -> thm                    \hfill{\bf infix}

RLN : thm list * (int * thm list) -> thm list  \hfill{\bf infix}

RL  : thm list * thm list -> thm list          \hfill{\bf infix}

MRL : thm list list * thm list -> thm list     \hfill{\bf infix}

\end{ttbox}

Joining rules together is a simple way of deriving new rules.  These

functions are especially useful with destruction rules.  To store

the result in the theorem database, use \ttindex{bind_thm}

(\S\ref{ExtractingAndStoringTheProvedTheorem}). 

\begin{ttdescription}

\item[\tt$thm@1$ RSN $(i,thm@2)$] \indexbold{*RSN} 

  resolves the conclusion of $thm@1$ with the $i$th premise of~$thm@2$.

  Unless there is precisely one resolvent it raises exception

  \xdx{THM}; in that case, use {\tt RLN}.

\item[\tt$thm@1$ RS $thm@2$] \indexbold{*RS} 

abbreviates \hbox{\tt$thm@1$ RSN $(1,thm@2)$}.  Thus, it resolves the

conclusion of $thm@1$ with the first premise of~$thm@2$.

\item[\tt {$[thm@1,\ldots,thm@n]$} MRS $thm$] \indexbold{*MRS} 

  uses {\tt RSN} to resolve $thm@i$ against premise~$i$ of $thm$, for

  $i=n$, \ldots,~1.  This applies $thm@n$, \ldots, $thm@1$ to the first $n$

  premises of $thm$.  Because the theorems are used from right to left, it

  does not matter if the $thm@i$ create new premises.  {\tt MRS} is useful

  for expressing proof trees.

\item[\tt$thms@1$ RLN $(i,thms@2)$] \indexbold{*RLN} 

  joins lists of theorems.  For every $thm@1$ in $thms@1$ and $thm@2$ in

  $thms@2$, it resolves the conclusion of $thm@1$ with the $i$th premise

  of~$thm@2$, accumulating the results. 

\item[\tt$thms@1$ RL $thms@2$] \indexbold{*RL} 

abbreviates \hbox{\tt$thms@1$ RLN $(1,thms@2)$}. 

\item[\tt {$[thms@1,\ldots,thms@n]$} MRL $thms$] \indexbold{*MRL} 

is analogous to {\tt MRS}, but combines theorem lists rather than theorems.

It too is useful for expressing proof trees.

\end{ttdescription}

\subsection{Expanding definitions in theorems}

\index{meta-rewriting!in theorems}

\begin{ttbox} 

rewrite_rule       : thm list -> thm -> thm

rewrite_goals_rule : thm list -> thm -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{rewrite_rule} {\it defs} {\it thm}]  

unfolds the {\it defs} throughout the theorem~{\it thm}.

\item[\ttindexbold{rewrite_goals_rule} {\it defs} {\it thm}]  

unfolds the {\it defs} in the premises of~{\it thm}, but leaves the

conclusion unchanged.  This rule underlies \ttindex{rewrite_goals_tac}, but 

serves little purpose in forward proof.

\end{ttdescription}

\subsection{Instantiating a theorem}

\index{instantiation}

\begin{ttbox}

read_instantiate    :            (string * string) list -> thm -> thm

read_instantiate_sg : Sign.sg -> (string * string) list -> thm -> thm

cterm_instantiate   :              (cterm * cterm) list -> thm -> thm

\end{ttbox}

These meta-rules instantiate type and term unknowns in a theorem.  They are

occasionally useful.  They can prevent difficulties with higher-order

unification, and define specialized versions of rules.

\begin{ttdescription}

\item[\ttindexbold{read_instantiate} {\it insts} {\it thm}] 

processes the instantiations {\it insts} and instantiates the rule~{\it

thm}.  The processing of instantiations is described

in \S\ref{res_inst_tac}, under {\tt res_inst_tac}.  

Use {\tt res_inst_tac}, not {\tt read_instantiate}, to instantiate a rule

and refine a particular subgoal.  The tactic allows instantiation by the

subgoal's parameters, and reads the instantiations using the signature

associated with the proof state.

Use {\tt read_instantiate_sg} below if {\it insts\/} appears to be treated

incorrectly.

\item[\ttindexbold{read_instantiate_sg} {\it sg} {\it insts} {\it thm}]

  resembles \hbox{\tt read_instantiate {\it insts} {\it thm}}, but reads

  the instantiations under signature~{\it sg}.  This is necessary to

  instantiate a rule from a general theory, such as first-order logic,

  using the notation of some specialized theory.  Use the function {\tt

    sign_of} to get a theory's signature.

\item[\ttindexbold{cterm_instantiate} {\it ctpairs} {\it thm}] 

is similar to {\tt read_instantiate}, but the instantiations are provided

as pairs of certified terms, not as strings to be read.

\end{ttdescription}

\subsection{Miscellaneous forward rules}\label{MiscellaneousForwardRules}

\index{theorems!standardizing}

\begin{ttbox} 

standard         :           thm -> thm

zero_var_indexes :           thm -> thm

make_elim        :           thm -> thm

rule_by_tactic   : tactic -> thm -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{standard} $thm$] puts $thm$ into the standard form

  of object-rules.  It discharges all meta-assumptions, replaces free

  variables by schematic variables, renames schematic variables to

  have subscript zero, also strips outer (meta) quantifiers and

  removes dangling sort hypotheses.

\item[\ttindexbold{zero_var_indexes} $thm$] 

makes all schematic variables have subscript zero, renaming them to avoid

clashes. 

\item[\ttindexbold{make_elim} $thm$] 

\index{rules!converting destruction to elimination}

converts $thm$, a destruction rule of the form $\List{P@1;\ldots;P@m}\Imp

Q$, to the elimination rule $\List{P@1; \ldots; P@m; Q\Imp R}\Imp R$.  This

is the basis for destruct-resolution: {\tt dresolve_tac}, etc.

\item[\ttindexbold{rule_by_tactic} {\it tac} {\it thm}] 

  applies {\it tac\/} to the {\it thm}, freezing its variables first, then

  yields the proof state returned by the tactic.  In typical usage, the

  {\it thm\/} represents an instance of a rule with several premises, some

  with contradictory assumptions (because of the instantiation).  The

  tactic proves those subgoals and does whatever else it can, and returns

  whatever is left.

\end{ttdescription}

\subsection{Taking a theorem apart}

\index{theorems!taking apart}

\index{flex-flex constraints}

\begin{ttbox} 

concl_of      : thm -> term

prems_of      : thm -> term list

nprems_of     : thm -> int

tpairs_of     : thm -> (term*term)list

stamps_of_thy : thm -> string ref list

theory_of_thm : thm -> theory

dest_state    : thm*int -> (term*term)list*term list*term*term

rep_thm       : thm -> {\ttlbrace}prop: term, hyps: term list, der: deriv, 

                        maxidx: int, sign: Sign.sg, shyps: sort list\ttrbrace

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{concl_of} $thm$] 

returns the conclusion of $thm$ as a term.

\item[\ttindexbold{prems_of} $thm$] 

returns the premises of $thm$ as a list of terms.

\item[\ttindexbold{nprems_of} $thm$] 

returns the number of premises in $thm$, and is equivalent to {\tt

  length(prems_of~$thm$)}.

\item[\ttindexbold{tpairs_of} $thm$] 

returns the flex-flex constraints of $thm$.

\item[\ttindexbold{stamps_of_thm} $thm$] 

returns the \rmindex{stamps} of the signature associated with~$thm$.

\item[\ttindexbold{theory_of_thm} $thm$]

returns the theory associated with $thm$.

\item[\ttindexbold{dest_state} $(thm,i)$] 

decomposes $thm$ as a tuple containing a list of flex-flex constraints, a

list of the subgoals~1 to~$i-1$, subgoal~$i$, and the rest of the theorem

(this will be an implication if there are more than $i$ subgoals).

\item[\ttindexbold{rep_thm} $thm$] decomposes $thm$ as a record containing the

  statement of~$thm$ ({\tt prop}), its list of meta-assumptions ({\tt hyps}),

  its derivation ({\tt der}), a bound on the maximum subscript of its

  unknowns ({\tt maxidx}), and its signature ({\tt sign}).  The {\tt shyps}

  field is discussed below.

\end{ttdescription}

\subsection{*Sort hypotheses} 

\index{sort hypotheses}

\begin{ttbox} 

force_strip_shyps : bool ref \hfill{\bf initially true}

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{force_strip_shyps}]

causes sort hypotheses to be deleted, printing a warning.

\end{ttdescription}

Isabelle's type variables are decorated with sorts, constraining them to

certain ranges of types.  This has little impact when sorts only serve for

syntactic classification of types --- for example, FOL distinguishes between

terms and other types.  But when type classes are introduced through axioms,

this may result in some sorts becoming {\em empty\/}: where one cannot exhibit

a type belonging to it because certain axioms are unsatisfiable.

If a theorem contains a type variable that is constrained by an empty

sort, then that theorem has no instances. It is basically an instance

of {\em ex falso quodlibet}.  But what if it is used to prove another

theorem that no longer involves that sort?  The latter theorem holds

only if under an additional non-emptiness assumption.

Therefore, Isabelle's theorems carry around sort hypotheses. The {\tt

shyps} field is a list of sorts occurring in type variables in the current

{\tt prop} and {\tt hyps} fields.  It may also includes sorts used in the

theorem's proof that no longer appear in the {\tt prop} or {\tt hyps}

fields --- so-called {\em dangling\/} sort constraints. These are the

critical ones, asserting non-emptiness of the corresponding sorts.

Isabelle tries to remove extraneous sorts from the {\tt shyps} field whenever

non-emptiness can be established by looking at the theorem's signature: from

the {\tt arities} information, etc.  Because its current implementation is

highly incomplete, the flag shown above is available.  Setting it to true (the

default) allows existing proofs to run.

\subsection{Tracing flags for unification}

\index{tracing!of unification}

\begin{ttbox} 

Unify.trace_simp   : bool ref \hfill{\bf initially false}

Unify.trace_types  : bool ref \hfill{\bf initially false}

Unify.trace_bound  : int ref \hfill{\bf initially 10}

Unify.search_bound : int ref \hfill{\bf initially 20}

\end{ttbox}

Tracing the search may be useful when higher-order unification behaves

unexpectedly.  Letting {\tt res_inst_tac} circumvent the problem is easier,

though.

\begin{ttdescription}

\item[Unify.trace_simp := true;] 

causes tracing of the simplification phase.

\item[Unify.trace_types := true;] 

generates warnings of incompleteness, when unification is not considering

all possible instantiations of type unknowns.

\item[Unify.trace_bound := $n$;] 

causes unification to print tracing information once it reaches depth~$n$.

Use $n=0$ for full tracing.  At the default value of~10, tracing

information is almost never printed.

\item[Unify.search_bound := $n$;] 

causes unification to limit its search to depth~$n$.  Because of this

bound, higher-order unification cannot return an infinite sequence, though

it can return a very long one.  The search rarely approaches the default

value of~20.  If the search is cut off, unification prints {\tt

***Unification bound exceeded}.

\end{ttdescription}

\section{Primitive meta-level inference rules}

\index{meta-rules|(}

These implement the meta-logic in {\sc lcf} style, as functions from theorems

to theorems.  They are, rarely, useful for deriving results in the pure

theory.  Mainly, they are included for completeness, and most users should

not bother with them.  The meta-rules raise exception \xdx{THM} to signal

malformed premises, incompatible signatures and similar errors.

\index{meta-assumptions}

The meta-logic uses natural deduction.  Each theorem may depend on

meta-level assumptions.  Certain rules, such as $({\Imp}I)$,

discharge assumptions; in most other rules, the conclusion depends on all

of the assumptions of the premises.  Formally, the system works with

assertions of the form

\[ \phi \quad [\phi@1,\ldots,\phi@n], \]

where $\phi@1$,~\ldots,~$\phi@n$ are the assumptions.  This can be

also read as a single conclusion sequent $\phi@1,\ldots,\phi@n \vdash

\phi$. Do not confuse meta-level assumptions with the object-level

assumptions in a subgoal, which are represented in the meta-logic

using~$\Imp$.

Each theorem has a signature.  Certified terms have a signature.  When a

rule takes several premises and certified terms, it merges the signatures

to make a signature for the conclusion.  This fails if the signatures are

incompatible. 

\index{meta-implication}

The {\bf implication} rules are $({\Imp}I)$

and $({\Imp}E)$:

\[ \infer[({\Imp}I)]{\phi\Imp \psi}{\infer*{\psi}{[\phi]}}  \qquad

   \infer[({\Imp}E)]{\psi}{\phi\Imp \psi & \phi}  \]

\index{meta-equality}

Equality of truth values means logical equivalence:

\[ \infer[({\equiv}I)]{\phi\equiv\psi}{\infer*{\psi}{[\phi]} &

                                       \infer*{\phi}{[\psi]}}  

   \qquad

   \infer[({\equiv}E)]{\psi}{\phi\equiv \psi & \phi}   \]

The {\bf equality} rules are reflexivity, symmetry, and transitivity:

\[ {a\equiv a}\,(refl)  \qquad

   \infer[(sym)]{b\equiv a}{a\equiv b}  \qquad

   \infer[(trans)]{a\equiv c}{a\equiv b & b\equiv c}   \]

\index{lambda calc@$\lambda$-calculus}

The $\lambda$-conversions are $\alpha$-conversion, $\beta$-conversion, and

extensionality:\footnote{$\alpha$-conversion holds if $y$ is not free

in~$a$; $(ext)$ holds if $x$ is not free in the assumptions, $f$, or~$g$.}

\[ {(\lambda x.a) \equiv (\lambda y.a[y/x])}    \qquad

   {((\lambda x.a)(b)) \equiv a[b/x]}           \qquad

   \infer[(ext)]{f\equiv g}{f(x) \equiv g(x)}   \]

The {\bf abstraction} and {\bf combination} rules let conversions be

applied to subterms:\footnote{Abstraction holds if $x$ is not free in the

assumptions.}

\[  \infer[(abs)]{(\lambda x.a) \equiv (\lambda x.b)}{a\equiv b}   \qquad

    \infer[(comb)]{f(a)\equiv g(b)}{f\equiv g & a\equiv b}   \]

\index{meta-quantifiers}

The {\bf universal quantification} rules are $(\Forall I)$ and $(\Forall

E)$:\footnote{$(\Forall I)$ holds if $x$ is not free in the assumptions.}

\[ \infer[(\Forall I)]{\Forall x.\phi}{\phi}        \qquad

   \infer[(\Forall E)]{\phi[b/x]}{\Forall x.\phi}   \]

\subsection{Assumption rule}

\index{meta-assumptions}

\begin{ttbox} 

assume: cterm -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{assume} $ct$] 

makes the theorem \(\phi \;[\phi]\), where $\phi$ is the value of~$ct$.

The rule checks that $ct$ has type $prop$ and contains no unknowns, which

are not allowed in assumptions.

\end{ttdescription}

\subsection{Implication rules}

\index{meta-implication}

\begin{ttbox} 

implies_intr      : cterm -> thm -> thm

implies_intr_list : cterm list -> thm -> thm

implies_intr_hyps : thm -> thm

implies_elim      : thm -> thm -> thm

implies_elim_list : thm -> thm list -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{implies_intr} $ct$ $thm$] 

is $({\Imp}I)$, where $ct$ is the assumption to discharge, say~$\phi$.  It

maps the premise~$\psi$ to the conclusion $\phi\Imp\psi$, removing all

occurrences of~$\phi$ from the assumptions.  The rule checks that $ct$ has

type $prop$. 

\item[\ttindexbold{implies_intr_list} $cts$ $thm$] 

applies $({\Imp}I)$ repeatedly, on every element of the list~$cts$.

\item[\ttindexbold{implies_intr_hyps} $thm$] 

applies $({\Imp}I)$ to discharge all the hypotheses (assumptions) of~$thm$.

It maps the premise $\phi \; [\phi@1,\ldots,\phi@n]$ to the conclusion

$\List{\phi@1,\ldots,\phi@n}\Imp\phi$.

\item[\ttindexbold{implies_elim} $thm@1$ $thm@2$] 

applies $({\Imp}E)$ to $thm@1$ and~$thm@2$.  It maps the premises $\phi\Imp

\psi$ and $\phi$ to the conclusion~$\psi$.

\item[\ttindexbold{implies_elim_list} $thm$ $thms$] 

applies $({\Imp}E)$ repeatedly to $thm$, using each element of~$thms$ in

turn.  It maps the premises $\List{\phi@1,\ldots,\phi@n}\Imp\psi$ and

$\phi@1$,\ldots,$\phi@n$ to the conclusion~$\psi$.

\end{ttdescription}

\subsection{Logical equivalence rules}

\index{meta-equality}

\begin{ttbox} 

equal_intr : thm -> thm -> thm 

equal_elim : thm -> thm -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{equal_intr} $thm@1$ $thm@2$] 

applies $({\equiv}I)$ to $thm@1$ and~$thm@2$.  It maps the premises~$\psi$

and~$\phi$ to the conclusion~$\phi\equiv\psi$; the assumptions are those of

the first premise with~$\phi$ removed, plus those of

the second premise with~$\psi$ removed.

\item[\ttindexbold{equal_elim} $thm@1$ $thm@2$] 

applies $({\equiv}E)$ to $thm@1$ and~$thm@2$.  It maps the premises

$\phi\equiv\psi$ and $\phi$ to the conclusion~$\psi$.

\end{ttdescription}

\subsection{Equality rules}

\index{meta-equality}

\begin{ttbox} 

reflexive  : cterm -> thm

symmetric  : thm -> thm

transitive : thm -> thm -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{reflexive} $ct$] 

makes the theorem \(ct\equiv ct\). 

\item[\ttindexbold{symmetric} $thm$] 

maps the premise $a\equiv b$ to the conclusion $b\equiv a$.

\item[\ttindexbold{transitive} $thm@1$ $thm@2$] 

maps the premises $a\equiv b$ and $b\equiv c$ to the conclusion~${a\equiv c}$.

\end{ttdescription}

\subsection{The $\lambda$-conversion rules}

\index{lambda calc@$\lambda$-calculus}

\begin{ttbox} 

beta_conversion : cterm -> thm

extensional     : thm -> thm

abstract_rule   : string -> cterm -> thm -> thm

combination     : thm -> thm -> thm

\end{ttbox} 

There is no rule for $\alpha$-conversion because Isabelle regards

$\alpha$-convertible theorems as equal.

\begin{ttdescription}

\item[\ttindexbold{beta_conversion} $ct$] 

makes the theorem $((\lambda x.a)(b)) \equiv a[b/x]$, where $ct$ is the

term $(\lambda x.a)(b)$.

\item[\ttindexbold{extensional} $thm$] 

maps the premise $f(x) \equiv g(x)$ to the conclusion $f\equiv g$.

Parameter~$x$ is taken from the premise.  It may be an unknown or a free

variable (provided it does not occur in the assumptions); it must not occur

in $f$ or~$g$.

\item[\ttindexbold{abstract_rule} $v$ $x$ $thm$] 

maps the premise $a\equiv b$ to the conclusion $(\lambda x.a) \equiv

(\lambda x.b)$, abstracting over all occurrences (if any!) of~$x$.

Parameter~$x$ is supplied as a cterm.  It may be an unknown or a free

variable (provided it does not occur in the assumptions).  In the

conclusion, the bound variable is named~$v$.

\item[\ttindexbold{combination} $thm@1$ $thm@2$] 

maps the premises $f\equiv g$ and $a\equiv b$ to the conclusion~$f(a)\equiv

g(b)$.

\end{ttdescription}

\subsection{Forall introduction rules}

\index{meta-quantifiers}

\begin{ttbox} 

forall_intr       : cterm      -> thm -> thm

forall_intr_list  : cterm list -> thm -> thm

forall_intr_frees :               thm -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{forall_intr} $x$ $thm$] 

applies $({\Forall}I)$, abstracting over all occurrences (if any!) of~$x$.

The rule maps the premise $\phi$ to the conclusion $\Forall x.\phi$.

Parameter~$x$ is supplied as a cterm.  It may be an unknown or a free

variable (provided it does not occur in the assumptions).

\item[\ttindexbold{forall_intr_list} $xs$ $thm$] 

applies $({\Forall}I)$ repeatedly, on every element of the list~$xs$.

\item[\ttindexbold{forall_intr_frees} $thm$] 

applies $({\Forall}I)$ repeatedly, generalizing over all the free variables

of the premise.

\end{ttdescription}

\subsection{Forall elimination rules}

\begin{ttbox} 

forall_elim       : cterm      -> thm -> thm

forall_elim_list  : cterm list -> thm -> thm

forall_elim_var   :        int -> thm -> thm

forall_elim_vars  :        int -> thm -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{forall_elim} $ct$ $thm$] 

applies $({\Forall}E)$, mapping the premise $\Forall x.\phi$ to the conclusion

$\phi[ct/x]$.  The rule checks that $ct$ and $x$ have the same type.

\item[\ttindexbold{forall_elim_list} $cts$ $thm$] 

applies $({\Forall}E)$ repeatedly, on every element of the list~$cts$.

\item[\ttindexbold{forall_elim_var} $k$ $thm$] 

applies $({\Forall}E)$, mapping the premise $\Forall x.\phi$ to the conclusion

$\phi[\Var{x@k}/x]$.  Thus, it replaces the outermost $\Forall$-bound

variable by an unknown having subscript~$k$.

\item[\ttindexbold{forall_elim_vars} $ks$ $thm$] 

applies {\tt forall_elim_var} repeatedly, for every element of the list~$ks$.

\end{ttdescription}

\subsection{Instantiation of unknowns}

\index{instantiation}

\begin{ttbox} 

instantiate: (indexname * ctyp){\thinspace}list * (cterm * cterm){\thinspace}list -> thm -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{instantiate} ($tyinsts$, $insts$) $thm$] 

simultaneously substitutes types for type unknowns (the

$tyinsts$) and terms for term unknowns (the $insts$).  Instantiations are

given as $(v,t)$ pairs, where $v$ is an unknown and $t$ is a term (of the

same type as $v$) or a type (of the same sort as~$v$).  All the unknowns

must be distinct.  The rule normalizes its conclusion.

\end{ttdescription}

\subsection{Freezing/thawing type unknowns}

\index{type unknowns!freezing/thawing of}

\begin{ttbox} 

freezeT: thm -> thm

varifyT: thm -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{freezeT} $thm$] 

converts all the type unknowns in $thm$ to free type variables.

\item[\ttindexbold{varifyT} $thm$] 

converts all the free type variables in $thm$ to type unknowns.

\end{ttdescription}

\section{Derived rules for goal-directed proof}

Most of these rules have the sole purpose of implementing particular

tactics.  There are few occasions for applying them directly to a theorem.

\subsection{Proof by assumption}

\index{meta-assumptions}

\begin{ttbox} 

assumption    : int -> thm -> thm Sequence.seq

eq_assumption : int -> thm -> thm

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{assumption} {\it i} $thm$] 

attempts to solve premise~$i$ of~$thm$ by assumption.

\item[\ttindexbold{eq_assumption}] 

is like {\tt assumption} but does not use unification.

\end{ttdescription}

\subsection{Resolution}

\index{resolution}

\begin{ttbox} 

biresolution : bool -> (bool*thm)list -> int -> thm

               -> thm Sequence.seq

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{biresolution} $match$ $rules$ $i$ $state$] 

performs bi-resolution on subgoal~$i$ of $state$, using the list of $\it

(flag,rule)$ pairs.  For each pair, it applies resolution if the flag

is~{\tt false} and elim-resolution if the flag is~{\tt true}.  If $match$

is~{\tt true}, the $state$ is not instantiated.

\end{ttdescription}

\subsection{Composition: resolution without lifting}

\index{resolution!without lifting}

\begin{ttbox}

compose   : thm * int * thm -> thm list

COMP      : thm * thm -> thm

bicompose : bool -> bool * thm * int -> int -> thm

            -> thm Sequence.seq

\end{ttbox}

In forward proof, a typical use of composition is to regard an assertion of

the form $\phi\Imp\psi$ as atomic.  Schematic variables are not renamed, so

beware of clashes!

\begin{ttdescription}

\item[\ttindexbold{compose} ($thm@1$, $i$, $thm@2$)] 

uses $thm@1$, regarded as an atomic formula, to solve premise~$i$

of~$thm@2$.  Let $thm@1$ and $thm@2$ be $\psi$ and $\List{\phi@1; \ldots;

\phi@n} \Imp \phi$.  For each $s$ that unifies~$\psi$ and $\phi@i$, the

result list contains the theorem

\[ (\List{\phi@1; \ldots; \phi@{i-1}; \phi@{i+1}; \ldots; \phi@n} \Imp \phi)s.

\]

\item[$thm@1$ \ttindexbold{COMP} $thm@2$] 

calls \hbox{\tt compose ($thm@1$, 1, $thm@2$)} and returns the result, if

unique; otherwise, it raises exception~\xdx{THM}\@.  It is

analogous to {\tt RS}\@.  

For example, suppose that $thm@1$ is $a=b\Imp b=a$, a symmetry rule, and

that $thm@2$ is $\List{P\Imp Q; \neg Q} \Imp\neg P$, which is the

principle of contrapositives.  Then the result would be the

derived rule $\neg(b=a)\Imp\neg(a=b)$.

\item[\ttindexbold{bicompose} $match$ ($flag$, $rule$, $m$) $i$ $state$]

refines subgoal~$i$ of $state$ using $rule$, without lifting.  The $rule$

is taken to have the form $\List{\psi@1; \ldots; \psi@m} \Imp \psi$, where

$\psi$ need not be atomic; thus $m$ determines the number of new

subgoals.  If $flag$ is {\tt true} then it performs elim-resolution --- it

solves the first premise of~$rule$ by assumption and deletes that

assumption.  If $match$ is~{\tt true}, the $state$ is not instantiated.

\end{ttdescription}

\subsection{Other meta-rules}

\begin{ttbox} 

trivial            : cterm -> thm

lift_rule          : (thm * int) -> thm -> thm

rename_params_rule : string list * int -> thm -> thm

rewrite_cterm      : thm list -> cterm -> thm

flexflex_rule      : thm -> thm Sequence.seq

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{trivial} $ct$] 

makes the theorem \(\phi\Imp\phi\), where $\phi$ is the value of~$ct$.

This is the initial state for a goal-directed proof of~$\phi$.  The rule

checks that $ct$ has type~$prop$.

\item[\ttindexbold{lift_rule} ($state$, $i$) $rule$] \index{lifting}

prepares $rule$ for resolution by lifting it over the parameters and

assumptions of subgoal~$i$ of~$state$.

\item[\ttindexbold{rename_params_rule} ({\it names}, {\it i}) $thm$] 

uses the $names$ to rename the parameters of premise~$i$ of $thm$.  The

names must be distinct.  If there are fewer names than parameters, then the

rule renames the innermost parameters and may modify the remaining ones to

ensure that all the parameters are distinct.

\index{parameters!renaming}

\item[\ttindexbold{rewrite_cterm} $defs$ $ct$]

transforms $ct$ to $ct'$ by repeatedly applying $defs$ as rewrite rules; it

returns the conclusion~$ct\equiv ct'$.  This underlies the meta-rewriting

tactics and rules.

\index{meta-rewriting!in terms}

\item[\ttindexbold{flexflex_rule} $thm$]  \index{flex-flex constraints}

removes all flex-flex pairs from $thm$ using the trivial unifier.

\end{ttdescription}

\index{meta-rules|)}

\section{Proof objects}\label{sec:proofObjects}

\index{proof objects|(} Isabelle can record the full meta-level proof of each

theorem.  The proof object contains all logical inferences in detail, while

omitting bookkeeping steps that have no logical meaning to an outside

observer.  Rewriting steps are recorded in similar detail as the output of

simplifier tracing.  The proof object can be inspected by a separate

proof-checker, or used to generate human-readable proof digests.

Full proof objects are large.  They multiply storage requirements by about

seven; attempts to build large logics (such as {\sc zf} and {\sc hol}) may

fail.  Isabelle normally builds minimal proof objects, which include only uses

of oracles.  You can also request an intermediate level of detail, containing

uses of oracles, axioms and theorems.  These smaller proof objects indicate a

theorem's dependencies.

Isabelle provides proof objects for the sake of transparency.  Their aim is to

increase your confidence in Isabelle.  They let you inspect proofs constructed

by the classical reasoner or simplifier, and inform you of all uses of

oracles.  Seldom will proof objects be given whole to an automatic

proof-checker: none has been written.  It is up to you to examine and

interpret them sensibly.  For example, when scrutinizing a theorem's

derivation for dependence upon some oracle or axiom, remember to scrutinize

all of its lemmas.  Their proofs are included in the main derivation, through

the {\tt Theorem} constructor.

Proof objects are expressed using a polymorphic type of variable-branching

trees.  Proof objects (formally known as {\em derivations\/}) are trees

labelled by rules, where {\tt rule} is a complicated datatype declared in the

file {\tt Pure/thm.ML}.

\begin{ttbox} 

datatype 'a mtree = Join of 'a * 'a mtree list;

datatype rule     = \(\ldots\);

type deriv        = rule mtree;

\end{ttbox}

%

Each theorem's derivation is stored as the {\tt der} field of its internal

record: 

\begin{ttbox} 

#der (rep_thm conjI);

{\out Join (Theorem "conjI", [Join (MinProof,[])]) : deriv}

\end{ttbox}

This proof object identifies a labelled theorem, {\tt conjI}, whose underlying

proof has not been recorded; all we have is {\tt MinProof}.

Nontrivial proof objects are unreadably large and complex.  Isabelle provides

several functions to help you inspect them informally.  These functions omit

the more obscure inferences and attempt to restructure the others into natural

formats, linear or tree-structured.

\begin{ttbox} 

keep_derivs  : deriv_kind ref

Deriv.size   : deriv -> int

Deriv.drop   : 'a mtree * int -> 'a mtree

Deriv.linear : deriv -> deriv list

Deriv.tree   : deriv -> Deriv.orule mtree

\end{ttbox}

\begin{ttdescription}

\item[\ttindexbold{keep_derivs} := MinDeriv $|$ ThmDeriv $|$ FullDeriv;] 

specifies one of the three options for keeping derivations.  They can be

minimal (oracles only), include theorems and axioms, or be full.

\item[\ttindexbold{Deriv.size} $der$] yields the size of a derivation,

  excluding lemmas.

\item[\ttindexbold{Deriv.drop} ($tree$,$n$)] returns the subtree $n$ levels

  down, always following the first child.  It is good for stripping off

  outer level inferences that are used to put a theorem into standard form.

\item[\ttindexbold{Deriv.linear} $der$] converts a derivation into a linear

  format, replacing the deep nesting by a list of rules.  Intuitively, this

  reveals the single-step Isabelle proof that is constructed internally by

  tactics.  

\item[\ttindexbold{Deriv.tree} $der$] converts a derivation into an

  object-level proof tree.  A resolution by an object-rule is converted to a

  tree node labelled by that rule.  Complications arise if the object-rule is

  itself derived in some way.  Nested resolutions are unravelled, but other

  operations on rules (such as rewriting) are left as-is.  

\end{ttdescription}

Functions {\tt Deriv.linear} and {\tt Deriv.tree} omit the proof of any named

theorems (constructor {\tt Theorem}) they encounter in a derivation.  Applying

them directly to the derivation of a named theorem is therefore pointless.

Use {\tt Deriv.drop} with argument~1 to skip over the initial {\tt Theorem}

constructor.

\index{proof objects|)}

\index{theorems|)}