Major Section: RULE-CLASSES

See rule-classes for a general discussion of rule classes and
how they are used to build rules from formulas. A `:built-in-clause`

rule can be built from any formula other than propositional
tautologies. Roughly speaking, the system uses the list of built-in
clauses as the first method of proof when attacking a new goal. Any
goal that is subsumed by a built in clause is proved ``silently.''

ACL2 maintains a set of ``built-in'' clauses that are used to
short-circuit certain theorem proving tasks. We discuss this at
length below. When a theorem is given the rule class
`:built-in-clause`

ACL2 flattens the `implies`

and `and`

structure of the
`:`

`corollary`

formula so as to obtain a set of formulas whose
conjunction is equivalent to the given corollary. It then converts
each of these to clausal form and adds each clause to the set of
built-in clauses.

For example, the following `:`

`corollary`

(regardless of the definition
of `abl`

)

(and (implies (and (true-listp x) (not (equal x nil))) (< (acl2-count (abl x)) (acl2-count x))) (implies (and (true-listp x) (not (equal nil x))) (< (acl2-count (abl x)) (acl2-count x))))will build in two clauses,

{(not (true-listp x)) (equal x nil) (< (acl2-count (abl x)) (acl2-count x))}and

{(not (true-listp x)) (equal nil x) (< (acl2-count (abl x)) (acl2-count x))}.We now give more background.

Recall that a clause is a set of terms, implicitly representing the
disjunction of the terms. Clause `c1`

is ``subsumed'' by clause `c2`

if
some instance of `c2`

is a subset `c1`

.

For example, let `c1`

be

{(not (consp l)) (equal a (car l)) (< (acl2-count (cdr l)) (acl2-count l))}.Then

`c1`

is subsumed by `c2`

, shown below,
{(not (consp x)) ; second term omitted here (< (acl2-count (cdr x)) (acl2-count x))}because we can instantiate

`x`

in `c2`

with `l`

to obtain a subset of
`c1`

.
Observe that `c1`

is the clausal form of

(implies (and (consp l) (not (equal a (car l)))) (< (acl2-count (cdr l)) (acl2-count l))),

`c2`

is the clausal form of
(implies (consp l) (< (acl2-count (cdr l)) (acl2-count l)))and the subsumption property just means that

`c1`

follows trivially
from `c2`

by instantiation.
The set of built-in clauses is just a set of known theorems in
clausal form. Any formula that is subsumed by a built-in clause is
thus a theorem. If the set of built-in theorems is reasonably
small, this little theorem prover is fast. ACL2 uses the ``built-in
clause check'' in four places: (1) at the top of the iteration in
the prover -- thus if a built-in clause is generated as a subgoal it
will be recognized when that goal is considered, (2) within the
simplifier so that no built-in clause is ever generated by
simplification, (3) as a filter on the clauses generated to prove
the termination of recursively `defun`

'd functions and (4) as a
filter on the clauses generated to verify the guards of a function.

The latter two uses are the ones that most often motivate an extension to the set of built-in clauses. Frequently a given formalization problem requires the definition of many functions which require virtually identical termination and/or guard proofs. These proofs can be short-circuited by extending the set of built-in clauses to contain the most general forms of the clauses generated by the definitional schemes in use.

The attentive user might have noticed that there are some recursive
schemes, e.g., recursion by `cdr`

after testing `consp`

, that ACL2 just
seems to ``know'' are ok, while for others it generates measure
clauses to prove. Actually, it always generates measure clauses but
then filters out any that pass the built-in clause check. When ACL2
is initialized, the clause justifying `cdr`

recursion after a `consp`

test is added to the set of built-in clauses. (That clause is `c2`

above.)

Note that only a subsumption check is made; no rewriting or
simplification is done. Thus, if we want the system to ``know''
that `cdr`

recursion is ok after a negative `atom`

test (which, by the
definition of `atom`

, is the same as a `consp`

test), we have to build
in a second clause. The subsumption algorithm does not ``know''
about commutative functions. Thus, for predictability, we have
built in commuted versions of each clause involving commutative
functions. For example, we build in both

{(not (integerp x)) (< 0 x) (= x 0) (< (acl2-count (+ -1 x)) (acl2-count x))}and the commuted version

{(not (integerp x)) (< 0 x) (= 0 x) (< (acl2-count (+ -1 x)) (acl2-count x))}so that the user need not worry whether to write

`(= x 0)`

or `(= 0 x)`

in definitions.
`:built-in-clause`

rules added by the user can be enabled and
disabled.

Major Section: RULE-CLASSES

See rule-classes for a general discussion of rule classes and
how they are used to build rules from formulas. An example
`:`

`corollary`

formula from which a `:compound-recognizer`

rule might be
built is:

Example: (implies (alistp x) When (alistp x) is assumed true, (true-listp x)) add the additional hypothesis that x is of primitive type true-listp.whereGeneral Forms: (implies (fn x) concl) (implies (not (fn x)) concl) (and (implies (fn x) concl1) (implies (not (fn x)) concl2)) (iff (fn x) concl) (equal (fn x) concl)

`fn`

is a Boolean valued function of one argument, `x`

is a
variable symbol, and the system can deduce some restriction on the
primitive type of `x`

from the assumption that `concl`

holds. The third
restriction is vague but one way to understand it is to weaken it a
little to ``and `concl`

is a non-tautological conjunction or
disjunction of the primitive type recognizers listed below.''The primitive ACL2 types and a suitable primitive recognizing expression for each are listed below.

type suitable primitive recognizerThus, a suitablezero (equal x 0) negative integers (and (integerp x) (< x 0)) positive integers (and (integerp x) (> x 0)) negative ratio (and (rationalp x) (not (integerp x)) (< x 0)) positive ratio (and (rationalp x) (not (integerp x)) (> x 0)) complex rational (complex-rationalp x) nil (equal x nil) t (equal x t) other symbols (and (symbolp x) (not (equal x nil)) (not (equal x t))) proper conses (and (consp x) (true-listp x)) improper conses (and (consp x) (not (true-listp x))) strings (stringp x) characters (characterp x)

`concl`

to recognize the naturals would be
`(or (equal x 0) (and (integerp x) (> x 0)))`

but it turns out that we
also permit `(and (integerp x) (>= x 0))`

. Similarly, the true-lists
could be specified by
(or (equal x nil) (and (consp x) (true-listp x)))but we in fact allow

`(true-listp x)`

. When time permits we will
document more fully what is allowed or implement a macro that
permits direct specification of the desired type in terms of the
primitives.
There are essentially four forms of `:compound-recognizer`

rules, the
last two general forms above being equivalent. We explain how such
rules are used by considering the individual forms.

Consider a rule of the first form, `(implies (fn x) concl)`

. The
effect of such a rule is that when the rewriter assumes `(fn x)`

true,
as it would while diving through `(if (fn x) xxx ...)`

to rewrite `xxx`

,
it restricts the type of `x`

as specified by `concl`

. However, when
assuming `(fn x)`

false, as necessary in `(if (fn x) ... xxx)`

, the rule
permits no additional assumptions about the type of `x`

. For example,
if `fn`

is `primep`

, i.e., the predicate that recognizes prime numbers,
then `(implies (primep x) (and (integerp x) (>= x 0)))`

is a compound
recognizer rule of the first form. When `(primep x)`

is assumed true,
the rewriter gains the additional information that `x`

is a natural
number. When `(primep x)`

is assumed false, no additional information
is gained -- since `x`

may be a non-prime natural or may not even be a
natural.

The second general form addresses itself to the symmetric case, when
assuming `(fn x)`

false permits type restrictions on `x`

but assuming
`(fn x)`

true permits no such restrictions. For example, if we
defined `exprp`

to be the recognizer for well-formed expressions for
some language in which all symbols, numbers, character objects and
strings were well-formed -- e.g., the well-formedness rules only put
restrictions on expressions represented by `consp`

s -- then the
theorem `(implies (not (exprp x)) (consp x))`

is a rule of the second
form. Assuming `(exprp x)`

true tells us nothing about the type of `x`

;
assuming it false tells us `x`

is a `consp`

.

The third general form addresses itself to the case where one type
may be deduced from `(fn x)`

and a generally unrelated type may be
deduced from its negation. If we modified the expression recognizer
above so that character objects are illegal, then a rule of the
third form is

(and (implies (exprp x) (not (characterp x))) (implies (not (exprp x)) (or (consp x) (characterp x)))).Finally, rules of the fourth class address the case where

`fn`

recognizes all and only the objects whose type is described. In
this case, `fn`

is really just a new name for some ``compound
recognizers.'' The classic example is `(booleanp x)`

, which is just a
handy combination of two primitive types:
(iff (booleanp x) (or (equal x t) (equal x nil))).

Often it is best to disable `fn`

after proving that it is a compound
recognizer, since `(fn x)`

will not be recognized as a compound
recognizer once it has been expanded.

Every time you prove a new compound recognizer rule about `fn`

it
overrides all previously proved compound recognizer rules about `fn`

.
Thus, if you want to establish the type implied by `(fn x)`

and you
want to establish the type implied by `(not (fn x))`

, you must prove a
compound recognizer rule of the third or fourth forms. Proving a
rule of the first form followed by one of the second only leaves the
second fact in the data base.

Compound recognizer rules can be disabled with the effect that older
rules about `fn`

, if any, are exposed.

If you prove multiple compound recognizer rules for a function, you
may see a **warning** message to the effect that the new rule is not as
``restrictive'' as the old. That is, the new rules do not give the
rewriter strictly more type information than it already had. The
new rule is stored anyway, overriding the old, if enabled. You may
be playing subtle games with enabling or rewriting. But two other
interpretions are more likely, we think. One is that you have
forgotten about an earlier rule and should merely print it out to
make sure it says what you know and then forget your new rule. The
other is that you meant to give the system more information and the
system has simply been unable to extract the intended type
information from the term you placed in the conclusion of the new
rule. Given our lack of specificity in saying how type information
is extracted from rules, you can hardly blame yourself for this
problem. Sorry. If you suspect you've been burned this way, you
should rephrase the new rule in terms of the primitive recognizing
expressions above and see if the warning is still given. It would
also be helpful to let us see your example so we can consider it as
we redesign this stuff.

Compound recognizer rules are similar to `:`

`forward-chaining`

rules in
that the system deduces new information from the act of assuming
something true or false. If a compound recognizer rule were stored
as a forward chaining rule it would have essentially the same effect
as described, when it has any effect at all. The important point is
that `:`

`forward-chaining`

rules, because of their more general and
expensive form, are used ``at the top level'' of the simplification
process: we forward chain from assumptions in the goal being proved.
But compound recognizer rules are built in at the bottom-most level
of the simplifier, where type reasoning is done.

All that said, compound recognizer rules are a rather fancy,
specialized mechanism. It may be more appropriate to create
`:`

`forward-chaining`

rules instead of `:compound-recognizer`

rules.

Major Section: RULE-CLASSES

See rule-classes for a general discussion of rule classes and
how they are used to build rules from formulas. An example
`:`

`corollary`

formula from which a `:congruence`

rule might be built is:

Example: (implies (set-equal x y) (iff (memb e x) (memb e y))).Also see defcong and see equivalence.

General Form: (implies (equiv1 xk xk-equiv) (equiv2 (fn x1... xk ...xn) (fn x1... xk-equiv ...xn)))where

`equiv1`

and `equiv2`

are known equivalence relations, `fn`

is an
`n-ary`

function symbol and the `xi`

and `xk-equiv`

are all distinct
variables. The effect of such a rule is to record that the
`equiv2`

-equivalence of `fn`

-expressions can be maintained if, while
rewriting the `kth`

argument position, `equiv1`

-equivalence is
maintained. See equivalence for a general discussion of the
issues. We say that `equiv2`

, above, is the ``outside equivalence''
in the rule and `equiv1`

is the ``inside equivalence for the `k`

th
argument''
The macro form `(defcong equiv1 equiv2 (fn x1 ... x1) k)`

is an
abbreviation for a `defthm`

of rule-class `:congruence`

that attempts to
establish that `equiv2`

is maintained by maintaining `equiv1`

in `fn`

's
`k`

th argument. The `defcong`

macro automatically generates the general
formula shown above. See defcong.

The `memb`

example above tells us that `(memb e x)`

is propositionally
equivalent to `(memb e y)`

, provided `x`

and `y`

are `set-equal`

. The
outside equivalence is `iff`

and the inside equivalence for the second
argument is `set-equal`

. If we see a `memb`

expression in a
propositional context, e.g., as a literal of a clause or test of an
`if`

(but not, for example, as an argument to `cons`

), we can rewrite
its second argument maintaining `set-equality`

. For example, a rule
stating the commutativity of `append`

(modulo set-equality) could be
applied in this context. Since equality is a refinement of all
equivalence relations, all equality rules are always available.
See refinement.

All known `:congruence`

rules about a given outside equivalence and `fn`

can be used independently. That is, consider two `:congruence`

rules
with the same outside equivalence, `equiv`

, and about the same
function `fn`

. Suppose one says that `equiv1`

is the inside equivalence
for the first argument and the other says `equiv2`

is the inside
equivalence for the second argument. Then `(fn a b)`

is `equiv`

`(fn a' b')`

provided `a`

is `equiv1`

to `a'`

and `b`

is `equiv2`

to `b'`

. This is an easy consequence of the transitivity of
`equiv`

. It permits you to think independently about the inside
equivalences.

Furthermore, it is possible that more than one inside equivalence
for a given argument slot will maintain a given outside equivalence.
For example, `(length a)`

is equal to `(length a')`

if `a`

and `a'`

are
related either by `list-equal`

or by `string-equal`

. You may prove two
(or more) `:congruence`

rules for the same slot of a function. The
result is that the system uses a new, ``generated'' equivalence
relation for that slot with the result that rules of both (or all)
kinds are available while rewriting.

`:congruence`

rules can be disabled. For example, if you have two
different inside equivalences for a given argument position and you
find that the `:`

`rewrite`

rules for one are unexpectedly preventing the
application of the desired rule, you can disable the rule that
introduced the unwanted inside equivalence.

More will be written about this as we develop the techniques.

Major Section: RULE-CLASSES

Example: (implies (true-listp x) (equal (len x) (if (null x) 0 (if (null (cdr x)) 1 (+ 2 (len (cddr x)))))))whereGeneral Form: (implies hyp (equiv (fn a1 ... an) body))

`equiv`

is an equivalence relation and `fn`

is a function
symbol other than `if`

, `hide`

, `force`

or `case-split`

. Such
rules allow ``alternative'' definitions of `fn`

to be proved as
theorems but used as definitions. These rules are not true
``definitions'' in the sense that they (a) cannot introduce new
function symbols and (b) do not have to be terminating recursion
schemes. They are just conditional rewrite rules that are
controlled the same way we control recursive definitions. We call
these ``definition rules'' or ``generalized definitions''.
Consider the general form above. Generalized definitions are stored
among the `:`

`rewrite`

rules for the function ``defined,'' `fn`

above, but
the procedure for applying them is a little different. During
rewriting, instances of `(fn a1 ... an)`

are replaced by corresponding
instances of `body`

provided the `hyp`

s can be established as for a
`:`

`rewrite`

rule and the result of rewriting `body`

satisfies the
criteria for function expansion. There are two primary criteria,
either of which permits expansion. The first is that the
``recursive'' calls of `fn`

in the rewritten body have arguments that
already occur in the goal conjecture. The second is that the
``controlling'' arguments to `fn`

are simpler in the rewritten body.

The notions of ``recursive call'' and ``controllers'' are complicated by the provisions for mutually recursive definitions. Consider a ``clique'' of mutually recursive definitions. Then a ``recursive call'' is a call to any function defined in the clique and an argument is a ``controller'' if it is involved in the measure that decreases in all recursive calls. These notions are precisely defined by the definitional principle and do not necessarily make sense in the context of generalized definitional equations as implemented here.

But because the heuristics governing the use of generalized
definitions require these notions, it is generally up to the user to
specify which calls in body are to be considered recursive and what
the controlling arguments are. This information is specified in the
`:clique`

and `:controller-alist`

fields of the `:definition`

rule class.

The `:clique`

field is the list of function symbols to be considered
recursive calls of `fn`

. In the case of a non-recursive definition,
the `:clique`

field is empty; in a singly recursive definition, it
should consist of the singleton list containing `fn`

; otherwise it
should be a list of all of the functions in the mutually recursive
clique with this definition of `fn`

.

If the `:clique`

field is not provided it defaults to `nil`

if `fn`

does
not occur as a function symbol in `body`

and it defaults to the
singleton list containing `fn`

otherwise. Thus, `:clique`

must be
supplied by the user only when the generalized definition rule is to
be treated as one of several in a mutually recursive clique.

The `:controller-alist`

is an alist that maps each function symbol in
the `:clique`

to a mask specifying which arguments are considered
controllers. The mask for a given member of the clique, `fn`

, must be
a list of `t`

's and `nil`

's of length equal to the arity of `fn`

. A `t`

should be in each argument position that is considered a
``controller'' of the recursion. For a function admitted under the
principle of definition, an argument controls the recursion if it is
one of the arguments measured in the termination argument for the
function. But in generalized definition rules, the user is free to
designate any subset of the arguments as controllers. Failure to
choose wisely may result in the ``infinite expansion'' of
definitional rules but cannot render ACL2 unsound since the rule
being misused is a theorem.

If the `:controller-alist`

is omitted it can sometimes be defaulted
automatically by the system. If the `:clique`

is `nil`

, the
`:controller-alist`

defaults to `nil`

. If the `:clique`

is a singleton
containing `fn`

, the `:controller-alist`

defaults to the controller
alist computed by `(defun fn args body)`

. If the `:clique`

contains
more than one function, the user must supply the `:controller-alist`

specifying the controllers for each function in the clique. This is
necessary since the system cannot determine and thus cannot analyze
the other definitional equations to be included in the clique.

For example, suppose `fn1`

and `fn2`

have been defined one way and it is
desired to make ``alternative'' mutually recursive definitions
available to the rewriter. Then one would prove two theorems and
store each as a `:definition`

rule. These two theorems would exhibit
equations ``defining'' `fn1`

and `fn2`

in terms of each other. No
provision is here made for exhibiting these two equations as a
system of equations. One is proved and then the other. It just so
happens that the user intends them to be treated as mutually
recursive definitions. To achieve this end, both `:definition`

rules
should specify the `:clique`

`(fn1 fn2)`

and should specify a suitable
`:controller-alist`

. If, for example, the new definition of `fn1`

is
controlled by its first argument and the new definition of `fn2`

is
controlled by its second and third (and they each take three
arguments) then a suitable `:controller-alist`

would be
`((fn1 t nil nil) (fn2 nil t t))`

. The order of the pairs in the
alist is unimportant, but there must be a pair for each function in
the clique.

Inappropriate heuristic advice via `:clique`

and `:controller-alist`

can
cause ``infinite expansion'' of generalized definitions, but cannot
render ACL2 unsound.

Note that the actual definition of `fn1`

has the runic name
`(:definition fn1)`

. The runic name of the alternative definition is
`(:definition lemma)`

, where `lemma`

is the name given to the event that
created the generalized `:definition`

rule. This allows theories to
switch between various ``definitions'' of the functions.

The definitional principle, `defun`

, actually adds `:definition`

rules. Thus the handling of generalized definitions is exactly the
same as for ``real'' definitions because no distinction is made in
the implementation. Suppose `(fn x y)`

is `defun`

'd to be
`body`

. Note that `defun`

(or `defuns`

or `mutual-recursion`

)
can compute the clique for `fn`

from the syntactic presentation and
it can compute the controllers from the termination analysis.
Provided the definition is admissible, `defun`

adds the
`:definition`

rule `(equal (fn x y) body)`

. Thus, `(fn a b)`

will not be expanded unless `(g a b)`

can be established.

Major Section: RULE-CLASSES

See rule-classes for a general discussion of rule classes and
how they are used to build rules from formulas. An example
`:`

`corollary`

formula from which an `:elim`

rule might be built is:

Example: (implies (consp x) when (car v) or (cdr v) appears (equal (cons (car x) (cdr x)) in a conjecture, and v is a x)) variable, consider replacing v by (cons a b), for two new variables a and b.whereGeneral Form: (implies hyp (equal lhs x))

`x`

is a variable symbol and `lhs`

contains one or more terms
(called ``destructor terms'') of the form `(fn v1 ... vn)`

, where `fn`

is a function symbol and the `vi`

are distinct variable symbols, `v1`

,
..., `vn`

include all the variable symbols in the formula, no `fn`

occurs in `lhs`

in more than one destructor term, and all occurrences
of `x`

in `lhs`

are inside destructor terms.
To use an `:elim`

rule, the theorem prover waits until a conjecture
has been maximally simplified. If it then finds an instance of some
destructor term `(fn v1 ... vn)`

in the conjecture, where the instance
for `x`

is some variable symbol, `vi`

, it instantiates the `:elim`

formula
as indicated by the destructor term matched, splits the conjecture
into two goals, according to whether the instantiated hypothesis,
`hyp`

, holds, and in the case that it does, generalizes all the
instantiated destructor terms in the conjecture to new variables and
then replaces `vi`

in the conjecture by the generalized instantiated
`lhs`

.

At the moment, the best description of how ACL2 `:elim`

rules are used
may be found in the discussion of ``ELIM Rules,'' pp 247 of A
Computational Logic Handbook.

Major Section: RULE-CLASSES

See rule-classes for a general discussion of rule classes and
how they are used to build rules from formulas. An example
`:`

`corollary`

formula from which a `:equivalence`

rule might be built is
as follows. (We assume that `r-equal`

has been defined.)

Example: (and (booleanp (r-equal x y)) (r-equal x x) (implies (r-equal x y) (r-equal y x)) (implies (and (r-equal x y) (r-equal y z)) (r-equal x z))).Also see defequiv.

General Form: (and (booleanp (equiv x y)) (equiv x x) (implies (equiv x y) (equiv y x)) (implies (and (equiv x y) (equiv y z)) (equiv x z)))except that the order of the conjuncts and terms and the choice of variable symbols is unimportant. The effect of such a rule is to identify

`equiv`

as an equivalence relation. Note that only Boolean
2-place function symbols can be treated as equivalence relations.
See congruence and see refinement for closely related
concepts.
The macro form `(defequiv equiv)`

is an abbreviation for a `defthm`

of
rule-class `:equivalence`

that establishes that `equiv`

is an
equivalence relation. It generates the formula shown above.
See defequiv.

When `equiv`

is marked as an equivalence relation, its reflexivity,
symmetry, and transitivity are built into the system in a deeper way
than via `:`

`rewrite`

rules. More importantly, after `equiv`

has been
shown to be an equivalence relation, lemmas about `equiv`

, e.g.,

(implies hyps (equiv lhs rhs)),when stored as

`:`

`rewrite`

rules, cause the system to rewrite certain
occurrences of (instances of) `lhs`

to (instances of) `rhs`

. Roughly
speaking, an occurrence of `lhs`

in the `kth`

argument of some
`fn`

-expression, `(fn ... lhs' ...)`

, can be rewritten to produce
`(fn ... rhs' ...)`

, provided the system ``knows'' that the value
of `fn`

is unaffected by `equiv`

-substitution in the `kth`

argument. Such knowledge is communicated to the system via
``congruence lemmas.''
For example, suppose that `r-equal`

is known to be an equivalence
relation. The `:`

`congruence`

lemma

(implies (r-equal s1 s2) (equal (fn s1 n) (fn s2 n)))informs the rewriter that, while rewriting the first argument of

`fn`

-expressions, it is permitted to use `r-equal`

rewrite-rules.
See congruence for details about `:`

`congruence`

lemmas.
Interestingly, congruence lemmas are automatically created when an
equivalence relation is stored, saying that either of the
equivalence relation's arguments may be replaced by an equivalent
argument. That is, if the equivalence relation is `fn`

, we store
congruence rules that state the following fact:
(implies (and (fn x1 y1) (fn x2 y2)) (iff (fn x1 x2) (fn y1 y2)))Another aspect of equivalence relations is that of ``refinement.'' We say

`equiv1`

``refines'' `equiv2`

iff `(equiv1 x y)`

implies
`(equiv2 x y)`

. `:`

`refinement`

rules permit you to establish such
connections between your equivalence relations. The value of
refinements is that if the system is trying to rewrite something
while maintaining `equiv2`

it is permitted to use as a `:`

`rewrite`

rule any refinement of `equiv2`

. Thus, if `equiv1`

is a
refinement of `equiv2`

and there are `equiv1`

rewrite-rules
available, they can be brought to bear while maintaining `equiv2`

.
See refinement.
The system initially has knowledge of two equivalence relations,
equality, denoted by the symbol `equal`

, and propositional
equivalence, denoted by `iff`

. `Equal`

is known to be a refinement of
all equivalence relations and to preserve equality across all
arguments of all functions.

Typically there are five steps involved in introducing and using a new equivalence relation, equiv.

More will be written about this as we develop the techniques. For now, here is an example that shows how to make use of equivalence relations in rewriting.(1) Define

`equiv`

,(2) prove the

`:equivalence`

lemma about`equiv`

,(3) prove the

`:`

`congruence`

lemmas that show where`equiv`

can be used to maintain known relations,(4) prove the

`:`

`refinement`

lemmas that relate`equiv`

to known relations other than equal, and(5) develop the theory of conditional

`:`

`rewrite`

rules that drive equiv rewriting.

Among the theorems proved below is

(defthm insert-sort-is-id (perm (insert-sort x) x))Here

`perm`

is defined as usual with `delete`

and is proved to be an
equivalence relation and to be a congruence relation for `cons`

and
`member`

.Then we prove the lemma

(defthm insert-is-cons (perm (insert a x) (cons a x)))which you must think of as you would

`(insert a x) = (cons a x)`

.
Now prove `(perm (insert-sort x) x)`

. The base case is trivial. The
induction step is

(consp x) & (perm (insert-sort (cdr x)) (cdr x))Opening-> (perm (insert-sort x) x).

`insert-sort`

makes the conclusion be
(perm (insert (car x) (insert-sort (cdr x))) x).Then apply the induction hypothesis (rewriting

`(insert-sort (cdr x))`

to `(cdr x)`

), to make the conclusion be
(perm (insert (car x) (cdr x)) x)Then apply

`insert-is-cons`

to get `(perm (cons (car x) (cdr x)) x)`

.
But we know that `(cons (car x) (cdr x))`

is `x`

, so we get `(perm x x)`

which is trivial, since `perm`

is an equivalence relation.Here are the events.

(encapsulate ((lt (x y) t)) (local (defun lt (x y) (declare (ignore x y)) nil)) (defthm lt-non-symmetric (implies (lt x y) (not (lt y x)))))(defun insert (x lst) (cond ((atom lst) (list x)) ((lt x (car lst)) (cons x lst)) (t (cons (car lst) (insert x (cdr lst))))))

(defun insert-sort (lst) (cond ((atom lst) nil) (t (insert (car lst) (insert-sort (cdr lst))))))

(defun del (x lst) (cond ((atom lst) nil) ((equal x (car lst)) (cdr lst)) (t (cons (car lst) (del x (cdr lst))))))

(defun mem (x lst) (cond ((atom lst) nil) ((equal x (car lst)) t) (t (mem x (cdr lst)))))

(defun perm (lst1 lst2) (cond ((atom lst1) (atom lst2)) ((mem (car lst1) lst2) (perm (cdr lst1) (del (car lst1) lst2))) (t nil)))

(defthm perm-reflexive (perm x x))

(defthm perm-cons (implies (mem a x) (equal (perm x (cons a y)) (perm (del a x) y))) :hints (("Goal" :induct (perm x y))))

(defthm perm-symmetric (implies (perm x y) (perm y x)))

(defthm mem-del (implies (mem a (del b x)) (mem a x)))

(defthm perm-mem (implies (and (perm x y) (mem a x)) (mem a y)))

(defthm mem-del2 (implies (and (mem a x) (not (equal a b))) (mem a (del b x))))

(defthm comm-del (equal (del a (del b x)) (del b (del a x))))

(defthm perm-del (implies (perm x y) (perm (del a x) (del a y))))

(defthm perm-transitive (implies (and (perm x y) (perm y z)) (perm x z)))

(defequiv perm)

(in-theory (disable perm perm-reflexive perm-symmetric perm-transitive))

(defcong perm perm (cons x y) 2)

(defcong perm iff (mem x y) 2)

(defthm atom-perm (implies (not (consp x)) (perm x nil)) :rule-classes :forward-chaining :hints (("Goal" :in-theory (enable perm))))

(defthm insert-is-cons (perm (insert a x) (cons a x)))

(defthm insert-sort-is-id (perm (insert-sort x) x))

(defun app (x y) (if (consp x) (cons (car x) (app (cdr x) y)) y))

(defun rev (x) (if (consp x) (app (rev (cdr x)) (list (car x))) nil))

(defcong perm perm (app x y) 2)

(defthm app-cons (perm (app a (cons b c)) (cons b (app a c))))

(defthm app-commutes (perm (app a b) (app b a)))

(defcong perm perm (app x y) 1 :hints (("Goal" :induct (app y x))))

(defthm rev-is-id (perm (rev x) x))

(defun == (x y) (if (consp x) (if (consp y) (and (equal (car x) (car y)) (== (cdr x) (cdr y))) nil) (not (consp y))))

(defthm ==-symmetric (== x x))

(defthm ==-reflexive (implies (== x y) (== y x)))

(defequiv ==)

(in-theory (disable ==-symmetric ==-reflexive))

(defcong == == (cons x y) 2)

(defcong == iff (consp x) 1)

(defcong == == (app x y) 2)

(defcong == == (app x y) 1)

(defthm rev-rev (== (rev (rev x)) x))