# Sentence Connectives in Formal Logic

*First published Tue May 4, 2010*

Natural language sentence-linking words (like “and”,
“or”) have traditionally had aspects of their use
simulated by corresponding connectives in various systems of formal
logic. (In fact we use the word “connective” more
generously, so that there don’t have to be two or more sentences
linked, allowing us to consider the analogues of words like
“not” and “necessarily” as connectives too.)
Amongst the questions this enterprise has raised are the
following. (1) Which such naturally occurring sentence connectives
call for a treatment of this kind? (Which of them should be counted
as ‘logical constants’?) (2) Do the aspects of their use
which are simulated in that process exhaust what we might think of as
their *meanings*, or is there typically some mismatch here? Of
course this second question depends on the particular formal treatment
given, since different logics may disagree on the formal principles
accorded to their representatives—for example in the case of
“if”, but according to one famous line of thought
(developed by H. P. Grice) when we focus specifically on their
treatment in *classical* propositional logic, any such apparent
discrepancies can be explained away by appeal to pragmatic principles
regulating ordinary cooperative conversation, and do not reflect
differences of meaning (semantic differences).

Neither question (1) nor question (2) will be addressed here.
Instead, this entry focusses on the properties of—and relations
between—the formal representatives of these connectives.
Examples of such properties include truth-functionality and a
replacement property called congruentiality, and of such relations what
we call the *subconnective* relation—relating one
connective to another when (roughly speaking) does every logical
principle holding for the first of a pair of connectives also holds for
the second. In fact, all of these concepts are subject to a further
relativity. For instance a connective is or is not congruential
*according to* a given logic, with a similar relativity in the
case of whether one connective is a subconnective of another. On the
other hand, as truth-functionality is defined below, a connective is
or is not truth-functional *with respect to* a class of
truth-value assignments—and various interesting further
relations give rise to a corresponding logic-relative version of this
concept depending on the different ways a class of truth-value
assignments is linked to a logic (which in turn depends on how we
think of a logic—as an ordinary consequence relation, for
example, or as a ‘multiple conclusion’ consequence
relation: see further note 1 below). For some other properties, the further relativity is a bit
different. For example, whether a given connective is *uniquely
characterized*—in the sense of being governed by rules which
when reduplicated to govern a new connective render compounds formed
from the same component sentences using the original and the new
connective completely interchangeable—is relative not just to a
logic but to a particular proof system (set of rules systematizing the
logic).

Section 1 introduces the idea of formal sentential languages as
(absolutely free) algebras, their primitive connectives having the
status of the fundamental operations of these algebras, and explains
the notion of a consequence relation on such a language. At a more
general level, Galois connections are also introduced, and a simple
observation made concerning what we call binary relational
connections, to be applied in later sections. Section 2 introduces the
fundamental idea of a valuation for a language (a bivalent truth-value
assignment to its formulas). Truth-functionality and a generalization
thereof called pseudo-truth-functionality (of a connective with
respect to a class of valuations) are explained, as well as the idea
of sequents and sequent-to-sequent rules. We then move to a more
general conception of sequents as originally conceived, allowing a set
of formulas to play the conclusion role rather than just a single
formula, as well as to the associated idea of generalized consequence
relations, emphasizing Carnap's motivation for making this
generalization. That motivation can be summarised with reference to
one connective. We take negation as our example. Consider the
consequence relation determined by the class of all valuations on
which the usual negation truth-function connective is associated with
the negation connective. This can be regarded as the classical logic
of negation, at least if logics are identified with consequence
relations. Now, there are additional valuations consistent with this
consequence relation on which the negation truth-function
is *not* associated with the negation connective. (The
terminology of a valuation's being consistent with a consequence
relation is adapted from Dana Scott. A precise definition appears in
Section 2.) But when we consider instead generalized consequence
relations, the analogous discrepancy does not arise. Section 2
includes a statement of a result of D. Gabbay which isolates precisely
the truth-functions for which such discrepancies do not arise
(roughly: generalizations of the conjunction truth-function and
various pro jection functions), as well as of a result of
W. Rautenberg showing that despite this expressive weakness in the
apparatus of consequence relations, they are sufficiently
discriminating—without the need to pass to generalized
consequence relations—to distinguish the logical properties of
any two truth-functions (of the same arity).

Section 3 picks up the theme of sequent-to-sequent rules, concentrating on semantic analysis in the style of J. W. Garson of such rules, and also on the existence and uniqueness of connectives satisfying given rules (a topic introduced by H. Hiż). These last two are seen to be connected with what we call cut-inductivity and id-inductivity of specifically sequent-calculus rules, via, in the former case, a proposal identifying the existence of a connective (with prescribed logical behaviour, given the perspective provided by a logic not explicitly providing such a connective) with conservative extension (of the given logic by rules for the new connective detailing that behaviour). This proposal was made by N. D. Belnap; some problems for conservative extension as a sufficient condition for granting the existence or intelligibility of a new connective with such-and-such logical properties are noted. Section 4 collects several cases of interest in which one might wonder about the existence of connectives with various such properties. For instance (this being Example 3 out of a total of six) one might ask if it makes sense to have a 1-ary connective with the property that two applications of this connective are equivalent to a single application of (classical) negation. Section 5 assembles notes and references.

## 1. Preliminaries

The connectives of the formal languages of propositional logic are
typically expected to exhibit salient features of the
sentence-connecting devices of natural languages, and especially such
devices as are regarded as items of logical vocabulary (*and*,
*or*, *not*, *if–then*,…), though it will
be no part of our concern to inquire as to what their
‘logicality’ consists in. This is an issue which arises
for expressions and constructions such as quantifiers, the identity
predicate, and so on, and not just for sentence connectives; for this
issue the entry “Logical Constants” by John MacFarlane
(2005) should be consulted. The formal languages concerned may be
taken to agree in constructing their formulas ultimately on the basis
of the same countably infinite supply of propositional
variables—or ‘sentence letters’—*p*_{1},
*p*_{2},…, the first three of which we write for
convenience as *p*, *q*, *r*. Where they differ is in respect of their stock
of primitive connectives. So for example, we have the language
**L** = (*L*, ∧, →, ¬, ⊥) with (set of)
formulas *L* and connectives as indicated, of arities 2, 2, 1
and 0. (We have no qualms in the last two cases in speaking of
connectives, even when there are not two or more formulas to
‘connect’. Note that 0-ary or nullary connectives are
already formulas—sometimes called sentential constants—in
their own right.) As the notation suggests (with boldface capitals for
algebras and the corresponding italic capitals for their universes),
we are thinking—in the Polish tradition—of languages as
algebras, identifying the primitive connectives with the fundamental
operations of such algebras, and in particular, as absolutely free
algebras of their similarity types (in the present instance, ⟨2,2,1,0⟩) with the *p*_{i} as free
generators. This allows us to speak in the usual fashion of the
*main connective* of a given formula, the variables
*occurring in* and more generally the *subformulas of*
such and such a formula, and so on, without restricting ourselves to
construing the formulas as obtained specifically by concatenating
various symbols.

Alongside this ‘syntactic operation’ sense
of *connective*—explained above for the primitive
connectives and extended to derived connectives
presently—another use of the term is indispensable: to refer to
a connective with particular logical behaviour, as when people want to
compare the behaviour of intuitionistic and classical implication. In
due course we will see the notion of a consequence relation as one way
of specifying such behaviour. If ⊢ is such a relation there is
a unique language in the above sense which is the language of ⊢
and we can think of this ‘logically loaded’ sense
of *connective* as applying to the ordered pair ⟨#,
⊢⟩; thus in the case of intuitionistic implication this
would be
⟨→, ⊢_{I L}⟩,
where ⊢_{I L} is
the consequence relation of intuitionistic logic. (In fact it would be
preferable to take equivalence classes of such pairs under a relation
of translational equivalence but we will make no further explicit use
of these ordered pairs or any such equivalence classes, and keep this
aspect of the discussion relatively informal. Instead of saying
“⟨#, ⊢⟩”, we will say things like
“# as it behaves according to ⊢”. It would be
better still to use not consequence relations—or the later
introduced generalized consequence relations—as the second
coordinate of the representatives ⟨#, ⊢⟩ but instead
to insert here a reference to a proof system—a set of
sequent-to-sequent rules of the kind introduced below, that
is—since important differences between logics fail to show up at
the level which would individuate them merely as consequence
relations.^{[1]})

The ‘languages as algebras’ perspective also allows us the convenience of defining matrix evaluations (see Section 5) as just homomorphisms from the algebra of formulas to the algebra of the matrix in question, of defining derived connectives and (sentential) contexts as term functions (formerly known as polynomial functions) and algebraic functions (now commonly called polynomial functions) respectively, and so on. This treatment would need some emendation were the discussion to embrace infinitary connectives, branching connectives, and various other exotic devices, but there will be enough to consider without touching on these.

For example, there is the issue raised by the use of the notation
above for the operations of **L**. Such a choice suggests that
∧, →, ¬ will represent conjunction, implication (or the
conditional) and negation—that is, be intended to capture some
aspect of the behaviour of *and*, *if–then*,
and *not*, respectively—and that ⊥ will be some kind
of falsity constant. (Similarly, while we are fixing notation, if we
had written “∨”, “↔”, or
“⊤”, we would be arousing expectations of
disjunction, equivalence—or biconditionality—and a truth
constant, respectively.) We can think of this as a matter of the
inference patterns to be sanctioned by some proof system offered for
**L**, or as a matter of how the formulas of this language are to
be interpreted by assigning suitable semantic values to them
(truth-values being the simplest candidates here). We refer to the
former as a proof-theoretic or syntactical characterization of the
logic and the latter as a semantical characterization. The terminology
is somewhat unfortunate in that a perfectly respectable view in the
philosophy of logic (going back at least to Gentzen) has it that
the *meanings* of items of logical vocabulary, such as the
connectives of current concern, for a particular reasoner, are given
by a specification of the set of primitive *rules*, or a
favoured subset thereof, governing those items in a proof system
representing that reasoner's (perhaps idealized) inferential
practices.^{[2]} We
take up some of those issues later, beginning with a reprise of the
semantical side of the story at the start of the following
section.

We need first some general preliminaries, and forgetting the intended
application(s) for a moment, let us define, given sets *S*
and *T*, a pair of functions (*f*, *g*) to be
a *Galois connection* between *S* and *T*
when
*f*: ℘(*S*) → ℘(*T*)
and
*g*: ℘(*T*)→ ℘(*S*),
and for all *S*_{0}, *S*_{1}
⊆ *S*, *T*_{0}, *T*_{1} ⊆ *T*, we have
*S*_{0} ⊆ *g*(*f*(*S*_{0})), *T*_{0} ⊆
*f*(*g*(*T*_{0})), *S*_{0} ⊆ *S*_{1} ⇒
*f*(*S*_{1}) ⊆ *f*(*S*_{0}), and
*T*_{0} ⊆
*T*_{1} ⇒ *g*(*T*_{1})
⊆ *g*(*T*_{0}). It is well known that the
composition of *g* with *f* (respectively, *f* with *g*)
in such a case is a *closure operation* on *S*
(respectively, on *T*) in the sense of being a function (we do
the *S* case)
*C*: ℘(*S*) → ℘(*S*) satisfying, for all
subsets *S*_{0}, *S*_{1}, of *S*:

S_{0}⊆C(S_{0})C(S_{0}) ⊆C(S_{0}∪S_{1})C(C(S_{0})) ⊆C(S_{0}).

Further, where *R* ⊆ *S* × *T*, the
triple ⟨*R*, *S*, *T*⟩, which we may
call a *binary relational connection* (between *S*
and *T*), induces a Galois connection
(*f*_{R},*g*_{R})
between *S* and *T*
when we put, for all *S*_{0}
⊆ *S*, *T*_{0}⊆ *T*:

f_{R}(S_{0}) = {t∈T: for alls∈S_{0}R(s,t)}

g_{R}(T_{0}) = {s∈S: for allt∈T_{0}R(s,t)}.

(“*S*” and “*T*” here are
mnemonic for “Source” and “Target”
respectively; note that we do not require *S* ∩ *T*
= ∅, or even that *S* ≠ *T*.) Suppose we
consider, not the sentential languages of our opening paragraph but
rather first-order languages, and take *S* as the set of closed
formulas of such a language and *T* as the set of models
(interpretations, structures,…) for that language, then with
*R* as the relation holding between a closed formula φ and a
structure M when
M ⊨ φ,
*f*_{R}(*S*_{0}) and *g*_{R}(*T*_{0}) are
usually denoted by *Mod*(*S*_{0})
and *Th*(*T*_{0}). (The labels are intended to suggest
“class of all models of” and “theory of”,
respectively.) In this case the derived closure operations
*Th* ○ *Mod*
(mapping *S*_{0}
to *Th*(*Mod*(*S*_{0}))) and
*Mod* ○ *Th*
(understood analogously) on *S* and *T* respectively deliver the set
of first-order consequences of *S*_{0} and the set of all
models verifying all the sentences verified by every
M ∈ *T*_{0}. When there is
only one such M in the latter case,
what is delivered is the class of structures elementarily equivalent
to M. (Here we are thinking of
classical first-order model theory.)

Returning to the former case: as this use of the term ‘consequences’
reminds us, a closure operation on the set of formulas of some
language is conventionally called a *consequence operation*,
and perhaps to use the notation *Cn*(⋅) rather than
simply *C*(⋅), though it is often more convenient to talk in terms
of the corresponding consequence *relation*, holding between a
set of formulas Γ and an individual formula φ—notated in some such way as “Γ ⊩ φ”, “Γ
⊢ φ”, etc.—when φ ∈ *Cn*(Γ);
when the relation has been defined semantically it is not uncommon to
press “⊨” into service in this role. (This will occur once
or twice below.) Certain shortcuts are usually taken with this
notation, so that one writes, for example “Γ, Δ, φ,
ψ ⊢ χ” and “⊢ φ” to mean “Γ ∪
Δ ∪ {φ,ψ} ⊢ χ” and “∅
⊢ φ”, respectively. (For the moment, capital Greek
letters range over sets of formulas.) The conditions given earlier for
*C*'s being a closure operation can be recast as conditions definitive
of a consequence relation, conditions we do not give explicitly
here. A consequence relation ⊢ on some language is
*finitary* if whenever Γ ⊢ φ then we have
Γ_{0} ⊢ φ for some finite Γ_{0} ⊆
Γ; to be *substitution-invariant* whenever Γ ⊢
φ then *s*(Γ) ⊢ *s*(φ) where *s* is any
substitution, i.e., any endomorphism of the algebra of formulas (and
*s*(Γ) is {*s*(ψ) | ψ ∈ Γ});
and, finally, to be *congruential* when φ ⊣⊢
ψ implies χ(φ) ⊣⊢ χ(ψ), in which the
notation is to be understood as follows. “φ ⊣⊢
ψ” abbreviates “φ ⊢ ψ and
ψ⊢ φ” (and if we are decorating the turnstile
“⊢” with superscripts or subscripts, these are
written only on the ⊢ and not also on the ⊣ part of the
composite notation); χ is some formula constructed from
propositional variables including at least
one, *p*_{k}, say—to indicate which we
may sometimes write χ = χ(*p*_{k})—and
χ(φ), for a formula φ, is the result of substituting φ
for all occurrences of *p*_{k} in χ. (So
χ(φ) is *s*(χ) for the unique
substitution *s* for which *s*(*p*_{i})
= *p*_{i} for *i* ≠ *k*
and *s*(*p*_{k}) = φ.) The
χ(⋅) notation amounts to thinking of χ as a
(sentential) *context*, and putting φ into this context
amounts to producing the formula χ(φ); more generally, we
can think of an *n*-ary context as a formula containing *n*
distinguished propositional variables
*p*_{k1},…, *p*_{kn},
say—and perhaps additional variables
too.^{[3]} If there
are no additional variables, then the formula
χ(*p*_{k1},…, *p*_{kn})
can be thought of as a derived or composite connective of the language
under consideration. Returning to the 1-ary case, when φ
⊣⊢ ψ we describe φ and ψ
as *equivalent* according to ⊢, and, following Smiley
(1962), when for all χ(⋅), the formulas χ(φ) and
χ(ψ) are equivalent according to ⊢, we describe
φ and ψ as
*synonymous* (according to ⊢). Thus a congruential
consequence relation is one for which any equivalent formulas are
synonymous.

The characterization of the (classical first-order) consequence
relation given above was semantic, in that we described it in terms of
the consequence
operation *Th* ○ *Mod* which invokes
the structures for the language and what is true in them. A proof
system for first-order logic provides a direct syntactic account of
when φ is a consequence of Γ, in terms of derivability using
rules (and perhaps axioms) sensitive only to the form of the formulas
involved. We could discuss the properties of sentence connectives in
terms of such quantified languages (in which open as well as closed
sentences can be their arguments) but to avoid extraneous
complications prefer the setting of the propositional languages with
which we began. And, as in the discussion to this point, we will
generally just say “connective”, rather than
“sentence connective”.

Any binary relational connection ⟨*R*, *S*, *T*⟩, as above,
will be said to have *conjunctive combinations* on the left
(right) if for any *s*_{0}, *s*_{1} ∈ *S* there exists *s*_{2} ∈ *S* such that
for all *t* ∈ *T*, *R*(*s*_{2}, *t*) just in case *R*(*s*_{0}, *t*) and
*R*(*s*_{1}, *t*) (respectively, for
any *t*_{0}, *t*_{1} ∈ *T* there exists *t*_{2} ∈ *T* such that for all *s* ∈
*S*, *R*(*s*, *t*_{2}) just in case *R*(*s*, *t*_{0}) and *R*(*s*, *t*_{1})); we call *s*_{2}
(respectively, *t*_{2}) a conjunctive combination of *s*_{0} and *s*_{1} (*t*_{0} and
*t*_{1}) in this case. Replacing the two “and”s after the words “just
in case” by “or” gives the notion of *disjunctive
combinations on* the left and right of
⟨*R*, *S*, *T*⟩. Both
notions admit of an obvious generalization to the case in which
{*s*_{0}, *s*_{1}} is replaced by an arbitrary subset *S*_{0} of *S*, and
similarly on the *T* side. We write ∏_{L} *S*_{0} and ∑_{L} *S*_{0} for conjunctive and disjunctive combinations, respectively, in
this case, and similarly on the right (with an “*R*” subscript), when
we know either that such a combination is unique or that, while not
unique, what is being said would be correct however a representative
combination were chosen. The subscript is omitted when no confusion is
likely. If, as above *S*_{0} = {*s*_{0}, *s*_{1}}, instead of writing ∏*S*_{0} (respectively, ∑*S*_{0}), we write *s*_{0} ⋅ *s*_{1} (respectively, *s*_{0} + *s*_{1}),
and similarly on the right. To illustrate these concepts, consider the
connection ⟨*R*, *S*, *T*⟩ with *S* some non-empty set, *T*
its power set, and *R* the relation of membership (∈). There are
conjunctive combinations and disjunctive combinations on the right,
being intersections and unions respectively, but there are in general
neither conjunctive nor disjunctive combinations on the left. Taking
the former case, a conjunctive combination of *s*_{0} and *s*_{1}, where
these are distinct elements of *S*, would be an element of *S*
belonging to precisely those subsets of *S* to which each of *s*_{0},
*s*_{1}, belongs—which is impossible since the unit set of this
element—as with any unit set—contains only one element, and so
not both of *s*_{0}, *s*_{1}. By analogy with the above definitions of
conjunctive and disjunctive combinations we can also consider the
existence of negative objects on the left and right of a relational
connection, and note that again the set-theoretic example just given
presents is with negative objects on the right (for *S*_{0} ⊆ *S*,
the corresponding negative object is *S* ∖ *S*_{0}), but not on
the left (“no negative individuals”, as philosophers might put it,
and the point would stand if we switched to the case in which *T*
comprised properties—rather than sets—of individuals,
reconstruing *R* accordingly). This absence of combinations
simultaneously on the left and right of a relational connection is
typical of non-degenerate cases, in that one can easily verify that if
a relational connection has conjunctive combinations on the left and
disjunctive combinations on the right, or vice versa, then it
satisfies the following (rather strong) *crossover
condition*:

for alls_{0},s_{1}∈S, allt_{0},t_{1}∈T, ifR(s_{0},t_{0}) andR(s_{1},t_{1}), then eitherR(s_{0},t_{1}) orR(s_{1},t_{0}).

(To see this, assume the antecedent and consider the relation *R* as
it relates *s*_{0} + *s*_{1} and *t*_{0} ⋅ *t*_{1}—or, for the “vice
versa”, *s*_{0} ⋅ *s*_{1} and *t*_{0} + *t*_{1}.)

## 2. Sequents and Valuations

Let us return to the propositional languages with which we began. A
*valuation* for such an **L** is a function from *L* to the
two-element set {T, F} (of truth-values). As in the case of
first-order languages touched on above, we put the set of formulas *L*
in the source position (i.e. on the left) of our relational
connections ⟨*L*, *V*, *R*⟩, with the set *V* of all
valuations for **L** as the target (on the right), and *R* (the
“is true on” relation) relates
φ ∈ *L* to
*v* ∈ *V* just in
case *v*(φ) = T. In view of the one-to-correspondence
between sets and characteristic functions, this relational connection
is just a mild variant of the ∈-based relational connection
considered above, in particular sharing with it the feature that it
has conjunctive and disjunctive combinations on the right but not the
left, and the feature that no two elements from the left bear the
relation *R* to exactly the same elements on the right, and no two
elements on the right have have exactly the same elements from the
left bearing the relation *R* to them. To get some interesting logic
under way, some of these features needed to be disturbed, though as we
shall see, one of them—the existence of conjunctive combinations on
the right—will prove surprisingly tenacious. If our interest lies
in classical logic, we will be especially keen to discard most of (the
above) *V* and concentrate only on boolean valuations, in the sense to
be explained now. A function from
{T, F}^{n} to {T, F} is called an *n*-ary *truth-function*, and we say
that an *n*-ary truth-function *f* *is associated with* the
*n*-ary connective # of **L** on the valuation *v* (for *L*)—or alternatively, under these same conditions, that # is associated
with *f* on *v*—just in case for all φ_{1},…,φ_{n}
∈ *L*:

v(#(φ_{1},…,φ_{n})) =f(v(φ_{1}),…,v(φ_{n})).

We describe the connective # as *truth-functional* with respect to a
class *U* of valuations when there is some *f* such that for each *v*
∈ *U*, *f* is associated with # on *v*, and as
*pseudo-truth-functional* with respect to *U* when *U* is a union of
classes of valuations with respect to each of which classes # is
truth-functional.^{[4]}
A valuation *v* is #-boolean, speaking
informally, when the expected truth-function is associated with # on
*v*. More specifically, *v*, assumed to be a valuation for a language
supporting the connectives to be mentioned here, is ∧-boolean
when the binary truth-function *f*_{∧} with *f*_{∧}(*x*, *y*) = T
iff *x* = *y* = T is associated with ∧ on *v*, ∨-boolean when
a similar association holds for ∨ and *f*_{∨} defined by
*f*_{∨}(*x*, *y*) = F iff
*x* = *y* = F. We assume that the idea is now
clear without providing separate definitions of
“→-boolean”, “↔-boolean”,
“¬-boolean”, “⊥-boolean”, etc. A
boolean valuation for **L** is just a valuation for **L**
which is #-boolean for each primitive connective # of **L** for
which a notion of #-boolean valuation has been defined. (This is just
a way of registering an intended truth-functional interpretation for
the connectives # in question.) Similarly we call any connective #
for which a notion of #-boolean valuation has been defined, a boolean
connective, even when discussing its behaviour according to a logic
not determined by a class of boolean valuations. (This terminology
will be explained in due course.)

A set of truth-functions is *strongly functionally complete* if
every truth-function can be obtained by composition from functions in
the set, together with the projection functions, and *weakly
functionally complete* if every truth-function can be obtained from
functions in the set, together with the projection functions and the
operation of passing from an (*n*+1)-ary function to an *n*-ary
function by dropping an inessential argument (an argument on which the
function does not depend, that is). This terminology is transferred
from truth-functions to boolean connectives via the conventional
association of such connectives with truth-functions described in the
previous paragraph. Thus for example each of the sets {∧,¬}, {→,¬} is weakly functionally complete, while
{→,⊥} is strongly functionally complete. This weak/strong
distinction (for more on which, see Humberstone (1993)) is not usually
registered in the literature, and functional completeness
*simpliciter* is understood as weak functional
completeness. These notions also apply in matrix semantics generally
(see Section 5), and there are
analogous notions of expressive completeness tailored to other areas;
two distinct versions of expressive completeness in the area of
intuitionistic logic (the distinction being orthogonal to the
weak/strong distinction just drawn) are contrasted in
Rousseau (1968), for example. None of these matters will
occupy us further here.

It is often useful to take a more fine-grained look at the
truth-functions in play in the above discussion. Assuming the
identification of *n*-ary functions with certain (*n*+1)-ary relations and
the latter with sets of ordered (*n*+1)-tuples, we use the term
*determinant* for the elements of a truth-function. Thus the
four determinants of the truth-function *f*_{→} associated with
→ on →-boolean valuations are ⟨T, T, T⟩, ⟨T, F, F⟩, ⟨F, T, T⟩, and ⟨F, F, T⟩. Note that the final entry is the ‘output’ value for the
function concerned. Such determinants will occupy us extensively
below.

In saying that a consequence relation ⊢ is *determined
by* a class *V* of valuations what is meant that Γ ⊢
ψ if and only if for all *v* ∈ *V*, whenever *v*(φ) = T
for each φ ∈ Γ, then *v*(ψ)= T. When ⊢
has been described by means of some proof system (some collection of
rules), then the “only if” (“if”) direction of
this claim amounts to the claim that ⊢ is *sound*
(respectively, *complete*) with respect to *V*. Take the language whose sole
primitive connective is ∧ and the class of ∧-boolean
valuations. A simple proof system is provided by the rules indicated
in the obvious way by the following figures:

(∧I)

φ ψ φ ∧ ψ

(∧E)

φ ∧ ψ φ ,

φ ∧ ψ ψ

The rule on the left is usually called ∧-Introduction while the
two on the right are collectively known as
∧-Elimination—thus “(∧I)” and
“(∧E)”—because they take us, respectively, to
and from a formula with ∧ as main connective, and they enable us
to define a consequence relation, ⊢_{∧} in purely syntactic
terms thus: Γ ⊢_{∧} φ just in case starting with
formulas all of which lie in Γ, successive application of the
above rules allows us to pass eventually to the formula φ. And
then one can show, not only that this is the consequence relation
determined by the class of ∧-boolean valuations—so that if
one began with that class and sought a syntactic codification of the
logic it gave rise to, one's search would be at an end—but
further that the valuations consistent with ⊢_{∧} are
precisely the ∧-boolean valuations. Here we calling
*v* *consistent with* ⊢ when for no Γ, ψ,
for which Γ ⊢ ψ, do we have *v*(φ) = T for all
φ ∈ Γ while *v*(ψ) = F; the “but
further” in the preceding sentence is explained by the fact that
every consequence relation ⊢ is determined by the class
Val(⊢) of all valuations consistent with
⊢. Typically ⊢ will also be determined by many further
classes of valuations, various subsets of Val(⊢), but
not so, as just remarked, in the present case. The significance of
this is that if one started not with the class of ∧-boolean
valuations (or any other kind of valuational semantics), but with the
rules (∧I) and (∧E) themselves, then that class of
valuations would force itself on one's attention anyway as the only
class of valuations, preservation of truth on each of which coincided
with derivability using those rules. (We return to this theme below,
in connection with Carnap.) Moreover the rules have a considerable
claim to the status of rules in accordance with which one must reason
if one is understanding ∧ as “and”, as providing a
codification of the way in which conjunctive premisses and conclusions
are reasoned from and to in natural day-to-day reasoning.

Now, as Gerhard Gentzen observed, when we come to formulate similarly
compelling principles for introducing and eliminating the other
connectives, we find ourselves having to make reference to the
*discharging* of assumptions temporarily made for the sake of
argument. Such is the case for instance with formulating an
introduction rule for → and with formulating an elimination rule
for ∨. In the case of the former rule, for example, Gentzen took it
that the way to derive φ → ψ is to assume φ and use
the various rules and other assumptions to derive ψ, at which
point the rule (also called Conditional Proof) allows us to claim that
we have derived φ → ψ from the remaining assumptions,
without having to record this conclusion as depending on φ (as
ψ itself did before the rule was applied). Thus to extend the
present approach, known as *natural deduction*, from ∧ to
these and other connectives, it is necessary to keep track not only of
what formulas we are dealing with at a given stage but also of the set
of initial formulas, or assumptions, on which that formula depends. We
want to say, in the case of the rule just described, that it entitles
us to pass from ψ, as depending on Γ ∪ {φ}, to
φ → ψ, as depending on Γ. The ordered pair of a
finite set of formulas and an individual formula, ⟨Γ,
φ⟩ whose first entry records the set of formulas on which its
second entry depends, we write in the more suggestive notation Γ
≻ φ, and call a *sequent*, or more explicitly, a
sequent of the logical framework Set-Fmla. For the
variant framework Set-Fmla_{0}, we allow also
the case in which there is no formula to the right of the
“≻”. Some principles for reasoning about sequents
which do not involve any specific connectives and were called by
Gentzen, *structural* rules (as opposed to
*operational* rules) are the following, which are respectively
a zero-premiss, a one-premiss and a two-premiss sequent-to-sequent
rule:

(Id) φ ≻ φ

(Weakening)

Γ ≻ ψ Γ,φ ≻ ψ

(Cut)

Γ,φ ≻ ψ Δ ≻ φ Γ, Δ ≻ ψ

As the notation makes clear, we use the same shortcuts in formulations
involving the regular turnstile (⊢) as for the current variant
(≻), despite the difference in role played by these two
symbols: the former is a metalinguistic predicate and the latter a
device for separating the two sides of a sequent. (“≻”,
arranging formulas into sequents, is, as Lemmon
(1965) says of his corresponding notation, the formal
analogue of the word “therefore”, which arranges sentences into
arguments. The former is, by contrast, for making comments about such
arguments—for example to the effect that they are valid. Of course
there is a little slack in the analogy for Set-Fmla, since
we allow the case where there are no premisses.) Any collection
Σ of sequents which is closed under the structural rules above
gives rise to a corresponding (finitary) consequence relation: Γ
⊢_{Σ} φ iff for some finite Γ_{0} ⊆ Γ,
we have Γ_{0}≻ φ ∈ Σ. We should also have a
sequent-based analogue of the notion of a valuation's being consistent
with a consequence relation; accordingly we say that a sequent Γ
≻ φ *holds on* the valuation *v* just in case we do
not have
*v*(ψ) = T for each ψ ∈ Γ while also
*v*(φ) = F. (When Γ is empty, the sequent holds on *v*
just in case *v*(φ) = T. It would be uncouth to speak of a
sequent as being true on a valuation, however, just as one would not
speak of the truth of an argument. Such usages are on a par with, and
closely allied to, the mistaken conception of ≻ as a kind of
connective.)

We can now formulate the desired →-introduction rule so that the assumption discharge feature is visible in what happens on the left of the sequent-separator as we move from premiss-sequent to conclusion-sequent. We write the elimination rule alongside in a similar notation; included also are the natural deduction rules for ∨; the earlier rules for ∧ should be imagined as re-written along similar lines, with assumption-dependencies made explicit:

(→I)

Γ,φ ≻ ψ Γ ≻ φ → ψ

(→E)

Γ ≻ φ → ψ Δ ≻ φ Γ, Δ ≻ ψ

(∨I)

Γ ≻ φ Γ ≻ φ ∨ ψ ,

Γ ≻ ψ Γ ≻ φ ∨ ψ

(∨E)

Γ ≻ φ ∨ ψ Δ,φ ≻ χ Θ,ψ ≻ χ Γ, Δ, Θ ≻ χ

Along with the ∧ rules re-formulated in the above sequent-to-sequent style, and the structural rule (Id)—(Weakening) and (Cut) being derivable—these rules render provable exactly those sequents in {∧,∨,→}-fragment of intuitionistic logic, while for the corresponding fragment of classical logic, one would need to add a further principle governing implication, such as the following subcontrariety rule:

(Subc _{→})

Γ, φ → ψ ≻ χ Γ, φ ≻ χ Γ ≻ χ

(Equivalently, in view of Weakening, one could give a formulation with
“Δ” as a set-variable for the right sequent
premiss, and carry both “Γ” and
“Δ” down to the conclusion-sequent, as in the
formulation of (→E) above.) In the terminology of
Dummett (1991) all of the rules we have seen
are *pure* (their schematic formulation involves mention only
of a single connective) and *simple* (each formulation mentions
it only once), this last rule is also *oblique*: a discharged
assumption has the connective in question as its main connective. One
can trade in this obliqueness for a loss of simplicity, changing the
→-elimination rule from the familiar Modus Ponens principle
above to a version incorporating Peirce's Law (“(*p* → *q*)
→ *p* ≻ *p*”), due to
Kanger:^{[5]}

Γ ≻ φ → ψ Δ ≻ (φ → χ) → φ Γ, Δ ≻ ψ

or sacrifice purity and have a zero-premiss rule ≻ φ ∨ (φ → ψ), or again just leave the implicational rules as they are and rely on further rules (not given here) governing negation to make up the shortfall and render such classically but not intuitionistically acceptable →-sequents—such as Peirce's Law above—provable with their aid. (The same tensions would then arise for the ¬ rules, if one regards obliqueness as undesirable, with “double negation elimination” principles abandoning simplicity, or “excluded middle” principles, abandoning purity. Note that by a ‘classically acceptable sequent’ here is meant a sequent holding on all boolean valuations.) An alternative for capturing the classical logic of → is to change from the logical framework Set-Fmla to Set-Set, in which finite (possibly empty) sets of formulas appear on the right of the ≻, just as on the left. Then the above →-Introduction rule can be remain intact, but with a set variable to collect side-formulas on the right: from Γ, φ≻ψ, Δ to Γ≻φ→ψ, Δ. (Suitable modifications should also be made to the remaining rules in this case.) These ‘multiple-succedent’ sequents are more popularly associated with a different approach to logic from the natural deduction approach (though they have often appeared in this setting), namely the sequent calculus approach, in which instead of being inserted or removed from the right of the ≻, by suitable introduction or elimination rules, together with the possible discharge of assumptions, the rules—again, ideally, pure and simple—always insert a connective, either to the left or the right, of the “≻”. The left insertion rules do the work of the elimination rules of the natural deduction approach, while the right insertion rules do the work of—indeed, are just—the natural deduction introduction rules. The left insertion rules, (# Left) for # ∈ {∧,∨,→} are then as follows for the classical case, with the right insertion rules (# Right) being the natural deduction introduction rules as above, with a set-variable on the right of the ≻ for the classical Set-Set case:

(∧ Left)

Γ,φ ≻ Δ Γ,φ ∧ ψ ≻ Δ ,

Γ,ψ ≻ Δ Γ,φ ∧ ψ ≻ Δ

(∨ Left)

Γ,φ ≻ Δ Γ′, ψ ≻ Δ′ Γ,Γ′,φ ∨ ψ ≻ Δ, Δ′

(→ Left)

Γ ≻ φ, Δ Γ′,ψ ≻ Δ′ Γ,Γ′,φ → ψ ≻ Δ, Δ′

For formulations restricted
to Set-Fmla_{0},
and thus suitable for the intuitionistic sequent calculus, we require
that Δ in (∧ Left) contains at most one formula, and for
(∨ Left) and (→ Left), that Δ ∪ Δ′
contains at most one formula. Indeed, one easily sees that using the
structural rules mentioned earlier and the operational rules described
here, there will always be exactly one formula, since proofs all begin
with (Id) and no rule removes all formulas from the right-hand side
(or ‘succedent’) in the transition from premiss sequents
to conclusion sequent. This changes if we add the rules for negation,
here giving both the left and the right insertion rules, initially for
the Set-Set
case:

(¬ Left)

Γ ≻ φ Δ Γ,¬φ ≻ Δ

(¬ Right)

Γ,φ ≻ Δ Γ ≻ ¬φ, Δ

A Set-Set sequent Γ≻ Δ holds on a
valuation *v* if and only if it is not the case that *v* verifies
every formula in Γ while falsifying every formula in Δ;
note that this coincides with the definition of “holding on *v*”
given for Set-Fmla in the case in which Δ =
{ψ} for some ψ. Similarly, we have the notion of a
*generalized consequence relation* in which suitably two-sided
versions of the conditions definitive of consequence relations are
given, which can be found in Shoesmith and Smiley
(1978), Chapter 2, for example (under the
name ‘multiple-conclusion consequence relation’). If
⊩ is such a relation on a language **L** then an
**L**-valuation *v* is said to be consistent with ⊩ when
there do not exist Γ, Δ, with
*v*(φ) = T for all
φ ∈ Γ and
*v*(ψ) = F for all
ψ ∈ Δ.
Sometimes formulations appear below with the phrase
“(generalized) consequence relation”, meaning that the
claim being made or considered applies both in the case of consequence
relations and in the case of generalized consequence relations.

One should distinguish sharply between *logical frameworks*,
that is, choices as to what kind of thing a sequent is
(e.g., Set-Fmla *vs.* Set-Set),
on the one hand, and *approaches to logic*, that is, choices as
to what kind of sequent-to-sequent rules are to be used (e.g., natural
deduction
*vs.* sequent calculus), on the other. Gentzen (1934)
used Set-Set
for a classical sequent calculus
and Set-Fmla
for the intuitionistic case. (More accurately, since he had finite
sequences of formulas in mind rather than sets, and in the latter case
allowed empty but not multiple succedents, we should really say
something long the lines
of Seq-Seq
and Seq-Fmla_{0},
respectively. Additional structural rules are then needed to allow
permutation of formulas in a sequence.) Rudolf Carnap (1943) had a rather different reason for passing
from Set-Fmla
to Set-Set,
one which is semantically based rather than grounded in considerations
of proof-theoretic elegance.

Let us forget about → and ¬ and the rules governing them,
and consider the formulas constructed with the aid only of ∧ and
∨. On sequents
of Set-Fmla
in this fragment, classical and intuitionistic logic agree, and the
consequence relation concerned—⊢_{∧,∨}, we may
call it—is the consequence relation determined by the class of
all ∧- and ∨-boolean valuations. Carnap's point is one
anticipated earlier with the remark that every *v*
∈ Val(⊢_{∧}) is ∧-boolean, which extends
to the case of ⊢_{∧,∨}: all valuations in
Val(⊢_{∧,∨}) are ∧-boolean, but they are
not all ∨-boolean, for the simple reason that for any consequence
relation ⊢, Val(⊢) is always closed under
conjunctive combinations, and it is easy to find ∨-boolean *u*,
*v*, with *u* ⋅ *v* not ∨-boolean.^{[6]} It is interesting to see this as an appeal
to the considerations of Section
1 as follows. Consider the binary relational connection
between the set of formulas in a language one of whose connectives is
∨ and the class of boolean valuations, with the relation
*is true on*. We have disjunctive combinations on the left,
with φ + ψ as
φ ∨ ψ, since the valuations on
the right are all ∨-boolean; therefore we cannot in general have
conjunctive combinations on the left, since the crossover condition
is not satisfied for this connection. The closure claim applies quite
generally, with ∏ *V* ∈ Val(⊢) whenever *V*
⊆ Val(⊢). This fact appears in the philosophical
literature in the fact that supervaluations over classes of bivalent
valuations deliver the same consequence relation as the original
valuations, which is accordingly classical when the valuations are
boolean. (Supervaluations are supposed to be non-bivalent, in
verifying those formulas verified by all the underlying valuations and
falsifying those formulas false on all the underlying valuations, with
those in neither category being neither true nor false on the
supervaluation. But everything in this definition after the words
“verified by all the underlying valuations” is mere
rhetoric, since all that matters—for a sequent's holding
on a valuation or a valuation's being consistent with a consequence
relation—is truth-preservation; setting it to one side, we can
identify the supervaluation over *V* with ∏ *V*. Likewise in the
case of ‘subvaluations’—as in Hyde
(2005)—where it is ∑ *V* that would be relevant.) A
special case arises with *V* = ∅, for which ∏ *V* is the
unique valuation *v*_{T} assigning T to every formula (of the language
in question, to which the claim of uniqueness is understood as
relativized).

One may wonder which truth-functions are like that associated (on
∧-boolean valuations) with ∧ in this respect. More
precisely, let ⊢ be the consequence relation determined by the
class of valuations on each of which an *n*-ary connective #_{f} is
associated with the *n*-ary truth-function *f*. Our question is: for
which choices of *f* does ⊢ enforce the interpretation of #_{f}
as *f* in the sense that on every *v* ∈ Val(⊢),
#_{f} is associated with *f*. This question was answered in Dov
Gabbay (1978) and (1981) (Chapter
1, Section 3) as follows. Say that *f* (as above) is a
*projection-conjunction* truth-function just in case for some
*J* ⊆ {1,…,*n*} we have, for all *x*_{1},…,*x*_{n} (*x*_{i}
∈ {T, F}):

f(x_{1},…,x_{n}) = T if and only if for allj∈J,x_{j}= T.

Then the truth-functions *f* for which every valuation consistent with
⊢ (as above) associates #_{f} with *f* are precisely the projection-conjunction
truth-functions. (Gabbay calls #_{f}, as it behaves according to
⊢, *strongly classical* in these cases, though his own
definition is not formulated in these terms.) By way of example: when
*n* = 1, these are the identity and the constant-true functions, while
for *n* = 2, they are the first projection, second projection,
constant true and (boolean) conjunction truth-functions. (The constant
true cases arise by taking *J* as ∅.) The remaining 12 binary
truth-functions thus exhibit the behaviour we saw with the disjunction
truth-function; the corresponding connectives are what Gabbay
calls *weakly* classical: determined by a class of valuations
over which the connective is associated with a certain truth-function
despite the existence of further consistent valuation on which it is
not so associated. It is easy to see that the latter phenomenon does
not arise in the case of generalized consequence relations ⊩,
since we can easily write down unconditional conditions on such
relations which enforce the association with a particular
truth-function on valuations in Val(⊩). This is best
seen as proceeding determinant by determinant. The unenforceable
determinant in the case of disjunction is ⟨F, F, F⟩,
dictating that when the disjuncts have the values listed in the first
two positions, the disjunction has the value listed in the final
position. This determinant induces in a familiar way (e.g., Segerberg
(1982), §3.3) the following condition on
⊩: that for all φ, ψ, we have φ ∨ ψ
⊩ φ, ψ. Alternatively put,
in Set-Set,
as opposed
to Set-Fmla,
there is a set of formulas holding on precisely the ∨-boolean
valuations, namely that containing, for all formulas φ, ψ,
the sequents φ ≻ φ ∨ ψ, and ψ ≻
φ ∨ ψ, and—the new case not already available
in Set-Fmla—φ
∨ ψ ≻ φ, ψ. In view of this, we can say that
in Set-Set
though not
in Set-Fmla,
the class of ∨-boolean valuations is *sequent-definable*,
and the same goes for connectives associated with all the other
truth-functions not meeting the projection-conjunction test (such as
negation and material implication). As the ∨ example makes
clear, by contrast with the case of consequence relations,
Val(⊩), for a generalized consequence relation
⊩, need not be closed under conjunctive
combinations.^{[7]}
Again, one can see this in the light of the considerations of
Section 1, thinking of the relational
connection
between Set-Set
sequents and arbitrary valuations, with the relation being
‘holds on’. This connection supplies disjunctive
combinations on the left, with σ + σ′ being the least
common weakening of σ and
σ′.^{[8]} Thus
it cannot also supply conjunctive combinations on the right, since it
does not satisfy the crossover condition.

This approximately Carnapian concern with sequent-definability—in particular the exclusion of unwanted non-boolean valuations—can be taken to a higher level by moving from questions about sequents and the valuations they hold, to questions about rules and the valuations they preserve the property of holding on, as we shall see in Section 3, after making some further points about the apparatus of consequence relations, or logics in the framework Set-Fmla. The determinant-induced conditions on generalized consequence relations alluded to above can be converted into conditions on consequence relations in a straightforward way, illustrated here in the case of negation, with the condition

for all φ: φ, ¬φ ⊩ ∅

(induced by ⟨T, F⟩) becoming

for all φ, ψ,: φ, ¬φ ⊢ ψ

and the condition

for all φ: ⊩ φ, ¬φ

(induced by ⟨F, T⟩) turning into the
*conditional* condition on consequence relations ⊩:

for all Γ, φ, ψ, if Γ, φ ⊢ ψ and Γ, ¬φ ⊢ ψ then Γ ⊢ ψ.

In a proof system
in Set-Fmla
such conditional requirements emerge as sequent-to-sequent rules,
closure under which (along with containing all the sequents
corresponding to the unconditional requirements) will suffice to
secure completeness—in this case that any unprovable sequent
fails to hold on some ¬-boolean valuation—without the
strong classicality property of having the provable sequents
hold *only* on such valuations. (The process just described
transforms
the Set-Set
condition,
φ ∨ ψ ≻ φ, ψ
(all φ, ψ), induced by the
Set-Fmla
unenforceable ∨ determinant
⟨F, F, F⟩, into
the *conditional* condition of closure under the rule
(∨E).) If a generalized consequence relation ⊩ with a
boolean connective # in its language satisfies all the conditions
induced by the determinants of the associated truth-function we call
⊩ #-*classical*. Similarly for a consequence relation
which satisfies the conditions on consequence
relations—*conditional* conditions
included—derived, as just described, from these
determinant-induced conditions on generalized consequence
relations. (These versions of #-classicality can often be simplified
on a case-by-case basis. For example, the ∧-classicality of a
consequence relation or generalized consequence relation—we use
the former notation here—is equivalent to its being the case
that for all φ, ψ, we have (1)
φ, ψ ⊢ φ ∧ ψ, (2)
φ ∧ ψ ⊢ φ, and (3)
φ ∧ ψ ⊢ ψ.
Here we see a three-part condition, instead of the four-part
condition comprising the conditions induced by each of the four
determinants.)

Consideration of the determinant-induced conditions on generalized
consequence relations allows us incidentally to observe, somewhat
surprisingly, that consequence relations (and
the Set-Fmla
framework) manage to distinguish the logics of any two truth-functions
of the same arity. More precisely, suppose that *f* and *g* are
distinct *n*-ary truth-functions, # is an *n*-ary connective, and
*V*_{f} and *V*_{g} are the classes of valuations on which, respectively,
*f* and *g* are associated with #; then the consequence relations
determined by *V*_{f} and *V*_{g} are distinct. By way of illustration,
suppose that *n* = 3 and *f* and *g* differ on some determinant
beginning with T, F, F. (Thus we are considering the case in which the
first component of a #-compound is true on a given valuation and the
second and third are false.) Without loss of generality, we can assume
that ⟨T, F, F, T⟩ ∈ *f* while ⟨T, F, F, F⟩
∈ *g*. The conditions on a generalized consequence relation
⊩ induced by these determinants are respectively
φ ⊩ ψ, χ, #(φ, ψ, χ)
and φ, #(φ, ψ, χ) ⊩ ψ, χ.
Thus in particular, taking φ as *p* and each of ψ,
χ, as *q*, for the generalized consequence relation
⊩_{f}, say, determined
by *V*_{f} and that,
⊩_{g}, determined
by *V*_{g}, we have:

p⊩_{f}q, #(p,q,q)and p, #(p,q,q) ⊩_{g}q.

Note that *p*, #(*p*, *q*, *q*) ⊮_{f} *q*, since otherwise it would
follow (by the ‘cut’ condition on generalized consequence relations)
from the first inset ⊩-statement here that *p* ⊩_{f} *q*,
which is not the case. Thus in view of the second ⊩-statement,
⊩_{f} and ⊩_{g} differ on the sequent σ = *p*, #(*p*, *q*,
*q*)≻ *q*, since σ ∈ ⊩_{g} while σ ∉
⊩_{f}. But σ is a sequent of Set-Fmla, so the
*consequence relations*, ⊢_{f} and ⊢_{g}, say,
determined respectively by *V*_{f} and *V*_{g} also differ (namely, on
σ). We have worked through a simple illustration of the general
result claimed above, but this case is sufficiently representative for
the general proof to be reconstructed.

Indeed the real moral is more general still. We did not need to use,
in our worked example, more than the fact that there was some
determinant respected by every valuation in *V*_{f} with the opposite
determinant respected by every valuation in *V*_{g}, where *V*
*respects* a determinant
⟨*x*_{1},…,*x*_{n},*x*_{n+1}⟩ for # (with each
*x*_{i} ∈ {T, F}) just in case for all formulas
φ_{1},…,φ_{n},
*v*(φ_{i}) = *x*_{i} for
*i* = 1,…, *n* implies
*v*(#(φ_{1},…,φ_{n}) = *x*_{n+1},
and determinants are
*opposite* if they differ in precisely their final
positions. (Actually, inspection of the worked example shows that we
also need the classes playing the role of *V*_{f} and *V*_{g} also to
contain enough valuations to assign different truth-values to any pair
of distinct propositional variables.) Take the 1-ary connective ☐
of any familiar modal logic (e.g., **S5**), for example. We
expect it to satisfy the determinant ⟨F, F⟩, in the sense that the appropriate consequence relation
should satisfy the condition induced by this determinant, since what
is necessarily true is true—and thus if φ is false,
☐φ should be false—while we do not want any determinant
of the form ⟨T, *x*⟩ to be satisfied. In this case
we can describe ☐ as *partially determined* according to
the consequence relation in question. On the other hand, if a
consequence relation ⊢ (or a generalized consequence relation
⊩) satisfies, for each *n*-tuple of truth-values, the condition
induced by exactly one determinant for the *n*-ary connective #,
whose first *n* entries are as given by that *n*-tuple (the final,
*n*+1^{st}, entry being either a T or an F), then we say # is
*fully determined* according to ⊢ (or ⊩). By the
determinant-induced conditions on a consequence relation here is meant
the conditions on a consequence relation derived as explained (or at
least illustrated) above from the conditions directly
determinant-induced conditions on a generalized consequence
relation. As one might expect, this property of connectives relative
to consequence relations has an intimate relation with the similar
property introduced at the start of this section:

A connective # in the language of a consequence relation ⊢ is fully determined according to ⊢ if and only if there is some class of valuationsVsuch that ⊢ is determined byVand # is truth-functional with respect toV.

Note that the two occurrences of the word ‘determined’
here mean quite different things. “Fully determined”, as
defined above, is a variation on Segerberg's terminology in
his (1982) of ‘type-determined’, varied so
as to lend itself to other modifications—*partially
determined*, *completely
undetermined*, *overdetermined*,… (though this last
will not be getting attention here)—while the second occurrence
of the word is the usual *sound and complete* sense of
“determined”. Turning to generalized consequence
relations, we find the connection to be more intimate still:

A connective # in the language of a generalized consequence relation ⊩ is fully determined according to ⊩ if and only if # is truth-functional with respect to Val(⊩).

This implies the statement given for consequence relations (putting
“⊩” in place of “⊢”) since a
generalized consequence relation, like a consequence relation, is
always determined by the class of all valuations consistent with
it. (But only the weaker statement above is available for consequence
relations, as such cases as that of ∨ illustrates, since the
consequence relation determined by the class of ∨-boolean
valuations—with respect to to which ∨ is by definition
truth-functional—is also determined by other classes of
valuations, with respect to it is not.) Indeed another way of putting the
contrast between the consequence relations claim above and the
generalized consequence relations claim is that while for consequence
relations #‘s being fully determined according to ⊢ is
equivalent *some* class of valuations which determines ⊢
having # truth-functional with respect to it, for generalized consequence
relations we can say that #‘s being fully determined according to
⊩ is equivalent *every* class of valuations which
determines ⊩ having # truth-functional with respect to it.

Along with truth-functionality, we also included a definition of
pseudo-truth-functionality with respect to *V*. A simple way of seeing the
relation between these two concepts is in an ∀/∃ scope
contrast (in which “*f*” ranges over truth-functions of the same arity
as #): # is truth-functional with respect to *V* when

∃f∀v∈V[fis associated with # onv],

whereas # is pseudo-truth-functional with respect to *V* when

∀v∈V∃f[fis associated with # onv].

It is natural to ask after a syntactical condition related to
pseudo-truth-functionality in the same way(s) as full determination is
related to truth-functionality. Say that an *n*-ary connective # is
*left-extensional* in its *i*^{th} position (1 ≤ *i* ≤
*n*) or *right-extensional* in its *i*^{th} position, according
to a generalized consequence relation ⊩, when the condition
(LE_{i}) or (RE_{i}) is satisfied, respectively, for all formulas
φ_{1},…,φ_{n},ψ in the language of ⊩ we
have:

LE _{i}: φ _{i}, ψ, #(φ_{1},…,φ_{n}) ⊩ #(φ_{1},…,φ_{i−1},ψ,φ_{i+1},…,φ_{n}),

RE _{i}: #(φ _{1},…,φ_{n}) ⊩ #(φ_{1},…,φ_{i−1},ψ,φ_{i+1},…,φ_{n}), φ_{i}, ψ.

# is left-extensional (right-extensional) according to ⊩ if it
is left-extensional (respectively, right-extensional) in its *i*^{th} position
for each *i*
(1 ≤ *i* ≤ *n*), and finally #
is *extensional* according to ⊩ when # is both left- and
right-extensional according to ⊩. For *n* = 1 we drop the
subscript, the above conditions reducing to (LE) and (RE):

LE : φ, ψ, #φ ⊩ #ψ,

RE : #φ⊩ #ψ, φ, ψ.

Thus (LE) requires that each *v* ∈
Val(⊩) “treats truths alike” when they are embedded
under #, in the sense that when
*v*(φ) = *v*(ψ) = T, then
*v*(#φ) = *v*(#ψ), while (RE) requires that falsehoods be
treated alike in that
*v*(φ) = *v*(ψ) = F likewise
implies that
*v*(#φ) = *v*(#ψ). (It may help to remark that for
↔-classical ⊢ together these amount to
requiring that
φ ↔ ψ ⊢ #φ ↔ #ψ.
But we are trying to keep the discussion
‘pure’ so that definitions make sense even when the connective—here
#—with the property being defined is the sole connective in the
language.) The conditions (LE_{i}) already make sense as conditions on
consequence relations (writing ⊢ for ⊩) since there is a
single formula on the right, and we will speak of # as
left-extensional according to ⊢ when they are all satisfied,
while the conditions (RE_{i}) need to be converted into conditional
conditions as above in the discussion of determinant-induced
conditions. We will illustrate this only for the *n* = 1 case, with
(RE) becoming (RE′), understood as demanding for all Γ,
φ, ψ, χ:

RE′ : If Γ, #ψ ⊢ χ and Γ , φ ⊢ χ and Γ , ψ ⊢ χ, then Γ, #φ ⊢ χ.

Taking it that the intention is clear for the general case what
(RE′_{i}) should be as a condition on consequence relations, we say
that # is right-extensional according to ⊢ when (RE′_{i}) is
satisfied for each *i* up to the arity of #, and that # is extensional
according to ⊢ just in case # is both left- and
right-extensional according to ⊢. Then paralleling the above
results for truth-functionality are the following for
pseudo-truth-functionality:

A connective # in the language of a consequence relation ⊢ is extensional according to ⊢ if and only if there is some class of valuationsVsuch that ⊢ is determined byVand # is pseudo-truth-functional with respect toV.

A connective # in the language of a generalized consequence relation ⊩ is extensional according to ⊩ if and only if # is pseudo-truth-functional with respect to Val(⊩).

Whether for consequence relations or for generalized consequence
relations, we have the following—in general, irreversible—implications for the three properties of being fully determined, being
extensional, and being congruential (as explained below) according to
any given (generalized) consequence relation:^{[9]}

fully determined⇒extensional⇒congruential.

where this last concept is defined by saying that # is
*congruential* in the *i*^{th} position according to a
consequence relation or generalized consequence relation ⊢ or
⊩ (we state the condition only using the “⊢” notation)
just in case for all formulas
φ_{1},…,φ_{n},ψ:

φ_{i}⊣⊢ ψ implies #(φ_{1},…,φ_{n}) ⊣⊢ #(φ_{1},…,φ_{i−1},ψ,φ_{i+1},…,φ_{n}),

and calling # congruential according to ⊢ when # is
congruential in each of its positions according to ⊢. Clearly a
consequence relation is congruential in the sense of Section
1—which of course also makes sense for generalized
consequence relations—precisely when each primitive connective in
its language is congruential in every position according to it. A
similar lifting of properties of connectives—‘according to’
(generalized) consequence relations—can be made in the other cases
too, for example calling a consequence relation fully determined when
each of the primitive connectives of its language is fully determined
according to it. Thus we may say that the consequence relation
⊢_{CL} of classical logic (with any given stock of primitive
connectives—though we will not further decorate the notation to
indicate a particular fragment) is fully determined and thus also
extensional (and congruential), while the consequence relation
⊢_{I L} of intuitionistic logic is congruential but not
extensional, since for example, ¬, is not right-extensional
according to ⊢_{I L}. The consequence relation determined by the
well-known three-element matrix of Łukasiewicz is not even
congruential, since (again) ¬ is not congruential according to
this relation. There is a weaker property than extensionality as here
defined (for the sake of the above connection with
pseudo-truth-functionality) which properly implies left extensionality
but not right extensionality and which, unlike left extensionality
itself, implies congruentiality. We could define #, here taken for
illustrative purposes again as 1-ary, to be weakly extensional
according to a consequence relation ⊢ when for all Γ,
φ, ψ:

Γ, φ ⊢ ψ and Γ, ψ ⊢ φ imply Γ, #φ ⊢ # ψ.

Thus weak extensionality is congruentiality ‘with
side formulas’ (collected into Γ). The obvious adaptation of
this notion to generalized consequence relations would be to allow
side-formulas on the right too (replacing the above“⊢” by
“⊩” and adding a set variable “Δ” to the right-hand
sides, that is): this turns out to be equivalent to extensionality for
generalized consequence relations as defined above in terms of
(LE_{i}) and (RE_{i}). ⊢_{I L} is weakly extensional in this
sense, and weak extensionality according to ⊢ ⊇
⊢_{I L} is a property of some theoretical significance for the
case in which the language of ⊢ adds new connectives in the
spirit of the original intuitionistic primitives. (For a more precise
articulation of what is at issue here, see Caicedo and Cignoli
(2001) and Caicedo (2004).)

Returning to the subject of extensionality, one area in which
extensional but not fully determined connectives arises is when one
takes hybrids of the latter. By the *hybrid* of disjunction and
conjunction, for example, understanding ∨ and ∧ to have the
inferential powers conferred on them by a particular logic, is meant a
binary connective whose inferential powers are those shared ∧
and ∨. (See Rautenberg (1989) for this
example, and Rautenberg (1991) and references
therein for more in this vein.) Note that the connectives hybridized
and the resulting hybrid are connectives in the ‘logically loaded’
sense of Section 1 above.

There are various ways of defining the hybridization process depend on
how much generality one wants, but the following minimally general
definition is the easiest to work with. Start with two (generalized)
consequence relations each on the same language, that language having
a single primitive connective, #, say. The hybrid connective is then
# as it behaves according to the intersection of the (generalized)
consequence relations. This way of defining things requires us to
re-notate the connectives to be hybridized. Thus sticking with the
example just mentioned, we consider a copy of the pure ∧
fragment of ⊢_{CL}, called ⊢_{∧} earlier in this
section, but writing # for ∧—but let us still call this
consequence relation ⊢_{∧}—and similarly with the ∨
fragment. We are then interested in the behaviour of # according to
⊢_{∧} ∩ ⊢_{∨}. (Where *V*_{∧} and *V*_{∨}
are the classes of valuations on which # is associated respectively
with the boolean ∧ truth-function and with the boolean ∨
truth-function, our intersective consequence relation is that
determined by *V*_{∧} ∪ *V*_{∨}.) This behaviour includes the
semilattice properties of idempotence, commutativity and
associativity, for example, since these are common to boolean
disjunction and conjunction. This particular hybrid connective is
partially determined, since for example
φ, ψ ≻ φ # ψ
(a sequent formulation of the condition induced on consequence
relations by ⟨T, T, T⟩) belongs to
⊢_{∧} ∩ ⊢_{∨}. This by itself does not show
that the hybrid is not *fully* determined, however, and to
avoid entanglement with some issues about rules, this point is best
illustrated in terms of the corresponding generalized consequence
relations and their intersection, the latter satisfying neither of the
conditions induced by the determinants beginning T, F: for ⟨T, F, T⟩, the condition that φ,
φ # ψ ⊩ ψ (failing for # as ∨) or the
determinant ⟨T, F, F⟩, for which the
induced condition is
φ ⊩ φ # ψ, ψ (failing for
# as ∧). (Amongst issues about non-zero-premiss
sequent-to-sequent rules that call for care is that fact that while
for any such rule under which each of the intersected consequence
relations is closed, their intersection is likewise closed, the
converse is false: the hybrid consequence relation may be closed under
rules for which the originals are not. In the present instance the
rule: *from* φ ≻ ψ *to* ψ ≻
φ is an example. Similarly for generalized consequence relations
the rule: *from* Γ ≻ Δ *to* Δ
≻ Γ makes the same point. This second version shows that
the current example of hybridization, as conducted in Set-Set, yields a symmetric generalized consequence
relation.^{[10]})

The hybrid of two fully determined connectives may even be completely
undetermined (satisfy no determinant-induced conditions at all, that
is). This will be the case, for example, if we hybridize not
conjunction and disjunction, but conjunction and negated conjunction
(sometimes called *nand*). By using negation in this way, we
prevent any overlap in the determinant-induced conditions which are
satisfied by the hybrid, though there still remain some non-trivial
properties (commutativity, for example) as well as extensionality—the point of this mention of hybrids—since any hybrid of fully
determined connectives will still be extensional (according to the
intersection, that is). Of course, one can also hybridize connectives
which are not themselves fully, or even partially determined—for
example, one could investigate the common properties of ☐ and
♢ according to the smallest normal modal logic (commonly
called **K**), though this example would for its most natural
development require us to go beyond the simple account of hybrids
given here, since it would have to attend to the simultaneous presence
of the boolean connectives alongside these modal operators. A better
known modal example of hybrids is that of the ☐ operator in
Łukasiewicz's Ł-modal logic (not that Łukasiewicz used this
notation for his necessity operator), which is a hybrid of 1-ary
connectives for the 1-place identity and constant-false
truth-functions. (See Chapter 1 of Prior (1957) on this
point; further references may be found in Section 5.)

The hybrid of two connectives is a *subconnective* of each of
them in the sense that whatever sequents (of the logical framework in
play) are provable for it are provable for them. And
conversely—but for the moment the subconnective relation itself
will be the focus of attention. Intuitionistic negation is in this
sense a subconnective, indeed a *proper* subconnective (in the
sense that we do not have the converse) of classical negation. As with
hybrids, there are various options, of differing degrees of generality
in their application, for defining this relation. In the case of a
language with only a single primitive connective, #, we can say that #
as it behaves according to (e.g.) the consequence relation ⊢ is
a subconnective of # as it behaves according to ⊢′ just in case
⊢ ⊆ ⊢′. If we want to allow different connectives of
the same arity, say # and #′, then we can put this instead by saying
that σ(#) ∈ ⊢ only if σ(#′) ∈ ⊢′,
for all σ(#), understanding the latter to denote a sequent with
various occurrences of # in its constituent formulas, with
σ(#′) the result of replacing all such occurrences by
#′. This kind of formulation lends itself to wider applicability
since we can allow other connectives in the formulas of σ(#),
provided these are in the language of ⊢′. There is no reason to
restrict attention to primitive connectives in making such
comparisons. For example, each of ☐ and ☐☐ is a
subconnective of the other in several well-known modal logics, as is
shown in Zolin (2000), where logics are developed axiomatically
as sets of formulas (sequents of the logical
framework Fmla, as we would
say^{[11]} so the
“σ(#)” would become “φ(#)” and
the point is that in the logics concerned φ(☐) and
φ(☐☐) are always equi-provable—though in the
interesting cases (e.g., those of **K**, **KT**, **KTB**), not
provably equivalent. Note that this is an intra-logical application of
the subconnective idea (corresponding to the case of ⊢ =
⊢′ above, for which we *definitely* need to work with
notationally distinguished # and #′). Zolin calls such mutual
subconnectives *analogous* (‘analogous
modalities’). A philosophically famous example of this
phenomenon from bimodal (or more specifically, doxastic-epistemic)
logic was reported in Byrd (1973).

Earlier in this section we saw that, where *V*_{f} and *V*_{g} are the
classes of valuations for a language with *n*-ary connective # on
which distinct truth-functions *f* and *g* are associated with #
respectively, the consequence relations determined by *V*_{f} and *V*_{g}
were always distinct. In the present terminology, this means that
either # as it behaves according to the first of these consequence
relations is not a subconnective of # as it behaves according to the
second, or vice versa. This is obvious for generalized consequence
relations as one can read off the difference from the unconditional
conditions induced by the determinants of *f* and *g*, but we saw (in
a worked example) that such differences are also guaranteed to emerge
at the relatively cruder level of consequence relations. A stronger
conclusion can be drawn from Theorem 4 of Rautenberg
(1981) (see also Rautenberg
(1985), p. 4), which will not be proved here,
and for the statement of which we need to explain that a
*non-trivial* consequence relation ⊢ is one for which
there exist formulas φ, ψ in the language of ⊢ such
that φ ⊬ ψ; for some of the matrix terminology here, see
Section 5:

LetMbe any two-element matrix with one designated value. Then the consequence relation ⊨_{M}determined byMis not properly included in any non-trivial substitution-invariant consequence relation on the same language.

As one very special case of this maximality result, in which the
algebra reduct of **M** has only a single fundamental operation
(*f*, as above, say), we get a stronger conclusion about the
consequence relations ⊢_{f} and ⊢_{g} determined by *V*_{f}
and *V*_{g} than our earlier conclusion that these consequence relations
are distinct: we get the conclusion that neither is included in the
other. Taking the **A** and *D* of **M** = (**A**, *D*) as
({T, F}, *f*) and {T}, the consequence relation
⊢_{f} is what Rautenberg is calling ⊨_{M}, and if
this were properly included in ⊢_{g} we should have a
contradiction with the above result, since the latter consequence
relation is certainly non-trivial and substitution-invariant. Thus
informally speaking, even in the relatively undiscriminating framework
Set-Fmla,
the logic of one truth-function never subsumes that of another. Less
informally: even
in Set-Fmla,
a fully determined connective is never a proper subconnective of
another fully determined connective.

We conclude this section with two observations prompted by this
result. The first concerns whether anything was gained by our earlier
looking at the argument for the weaker result (understood in the light
of the above conventions) to the effect that, *f* ≠ *g* implies
⊢_{f} ≠ ⊢_{g}. The answer to this question is that
although the result is weaker, the proof yields something stronger:
any difference in determinants yields a difference in unconditional
determinant-induced conditions. Considered as sequents (with
propositional variables replacing the schematic letters in the induced
conditions) these have a very special form, and, transferring
Dummett's terminology (seen earlier in this section) for rules, across
to sequents, we note that they are all included in the class of
*pure and simple* sequents: sequents in which there is only one
connective to occur and it occurs only once in the sequent. Suitably
abstracting from our earlier worked example, we can see that any
connectives satisfying the conditions induced by even one pair of
opposite determinants (whether or not any remaining
determinant-induced conditions are satisfied) are bound to differ on a
pair of pure and simple sequents. But the analogue of Rautenberg's
result does not hold in this form: that the set of pure and simple
sequents satisfied by the one and the set satisfied by the other are
incomparable with respect to inclusion. For example, consider the case of the
1-ary truth-functions boolean negation and the constant-false
function. We can cite #*p* ⊢ *q* as an example of a pure and
simple sequent which holds on every valuation on which 1-ary # is
associated with the constant-false function but not on every valuation
on which it is associated with the negation function. But an
inspection of possible cases reveals that there is no pure and simple
sequent valid for the negation interpretation of # which is not also
valid for the constant-false interpretation. (Of course there are
sequents doing this—e.g., *p* ≻ ## *p* but, though pure,
they are not simple. And, yes, in Set-Set, we can always
find pure and simple witnesses to the non-inclusion of the one logic
in the other: the determinant-induced conditions on generalized
consequence relations will serve in this capacity.)

For the second observation, we note a more linguistic reformulation of
Rautenberg's result, for which we first need a brief digression on
defined connectives. A connective derived by composition from a set of
primitive connectives yields a defined connective in one of two ways,
depending on one's views of definition. On one view, isolating the
derived connective is already giving a definition: one has a formula
φ(*p*_{1},…, *p*_{n}) in which exactly the exhibited variables occur,
which since it is constructed from these variables via the primitive
connectives constitutes a derived connective, applying to formulas
ψ_{1},…, ψ_{n} to yield the formula
φ(ψ_{1},…,ψ_{n})
which results from substituting
ψ_{i} uniformly for *p*_{i} in
φ(*p*_{1},…, *p*_{n}).
We can introduce a
shorthand to such substitution-instances, denoting
φ(ψ_{1},…,ψ_{n}) by
“#(ψ_{1},…,ψ_{n})”, for
instance—as in calling the formula
¬(ψ_{1} ∧ ¬ ψ_{2})
by the name “ψ_{1} → ψ_{2}”. (Of course this would only be a
good idea if there is nothing better to mean, or which is already
meant, by →—as there would be in intuitionistic logic, for
instance.) That is the first view of definition.^{[12]} On the second view, rather than just being a
metalinguistic abbreviation, the defined connective is an addition to
the object language: we add (to continue with the example just given)
a new primitive connective →, along with a stipulation that
compounds of the form ψ_{1} → ψ_{2} are to be synonymous with
those of the corresponding form ¬(ψ_{1} ∧ ¬ ψ_{2}). We
need not choose between these views here, but some formulations are
sensitive to which view is taken. In particular, if we have classical
logic with primitives ¬ and ∧, we may still wish to
consider its implicational fragment—in Set-Fmla, say,
where this would be the restriction of ⊢_{CL} to formulas
constructed using only the connective →. On the first view of
definition, we cannot literally do that, since to use → is
*ipso facto* to use ¬ and ∧, so what is meant is
rather than the only occurrences of ¬ and ∧ are such as
occur in subformulas of the form ψ_{1} → ψ_{2} in virtue of
these being the formulas ¬(ψ_{1} ∧ ¬ψ_{2}). On the
second view of definition, we have expanded the language-as-algebra by
adding a new fundamental operation → and can now consider the
subalgebra generated by the propositional variables of the reduct of
this algebra obtained by discarding the other fundamental operations.

Now let φ_{1} and φ_{2} any two classically non-equivalent
formulas in the same variables, *p*_{1},…, *p*_{n} for definiteness, and
consider the consequence relations ⊢_{1} and ⊢_{2}, where
⊢_{1} is the restriction of ⊢_{CL} to formulas constructed
using the *n*-ary connective # defined thus:

#(ψ_{1},…,ψ_{n}) = φ_{1}(ψ_{1},…,ψ_{n}),

and ⊢_{2} is the analogous restriction of ⊢_{CL} this time
to the #-fragment with # defined by:

#(ψ_{1},…,ψ_{n}) = φ_{2}(ψ_{1},…,ψ_{n}).

Since each of φ_{1} and φ_{2} fixes an *n*-ary truth-function
of the variables involved, and these truth-functions are distinct
(since φ_{1} and φ_{2} were to be non-equivalent) the above
*V*_{f} / *V*_{g} application of Rautenberg's result tells us that neither
of ⊢_{1}, ⊢_{2}, is included in the other. Concentrating on
the connectives involved, we can put this by saying that in classical
logic any two non-equivalent connectives are incomparable with respect
to the subconnective relation (understood as relating the consequence
relations on languages in which no further connectives appear). We
close by observing that the analogous claim would be false for
intuitionistic logic. A counterexample is provided by the connectives
defined by
φ #_{1} ψ = ¬φ∨ ¬ψ and
φ #_{2} ψ = ¬(φ ∧ ψ).
These are intuitionistically
non-equivalent, in the sense of not always yielding
⊢_{I L}-equivalent compounds from the same components,
while—by an argument not given here—for any
intuitionistically provable sequent σ(#_{1}) involving formulas
constructed from the variables by means of #_{1}, the corresponding
sequent σ(#_{2}) is also intuitionistically provable (though not
conversely). So, by contrast with what can happen in the classical
setting, we have non-equivalent connectives one of which is a proper
subconnective of the other—understanding this last phrase
subject to the same parenthetical qualification as above.

## 3. Rules and Connectives

For this section the focus will be on rules governing the various connectives, beginning with a deferred topic from Section 2: the transfer of Carnap-inspired questions of the range of valuations on which the sequents provable in this or that logic hold to question about the range of valuations on which the sequent-to-sequent rules preserve the property of holding. We will also revisit the sequent-calculus operational rules from early in Section 2 and look at some of their properties, as well as at some interesting variations on the structural rules which work alongside them. One topic—or pair of topics—raised by such investigations is the issue of the existence and uniqueness of connectives exhibiting prescribed logical behaviour (such as might be enshrined in various rules), an issue bearing on all approaches to logic and not just the sequent calculus approach. Finally we look at a pair of concepts applying to connectives as they behave in particular logics: the concept of a universally representative connective and the concept of a ‘special’ connective.

Beginning, then, with the preservation behaviour of rules (as in
Garson (1990) and (2001)) it is
necessary first to distinguish a (sequent-to-sequent) rule's having a
local preservation characteristic with respect to a class *V* of
valuations—that for each *v* ∈ *V*, the rule preserves the property
of holding on *v*—from its having a global preservation
characteristic with respect to *V*: if the premiss-sequents hold on
every *v* ∈ *V*, then the conclusion-sequent holds on every *v* ∈
*V*. For brevity, a sequent will be called *V*-*valid* when it
holds on each *v* ∈ *V*. Then we distinguish the local range of a rule
ρ, denoted *Loc*(ρ) from its global range,
*Glo*(ρ), taking the definitions as relative to a
background language the class of all valuations for which we denote by
*U*:

Loc(ρ) = {v∈U: ρ preserves the property of holding onv}

Glo(ρ) = {V⊆U: ρ preservesV-validity}.

It is also natural to define the local and global range of a
collection *R* of rules, by setting *Loc*(*R*) = ∩_{ρ
∈ R}*Loc*(ρ) and

Glo(R) = {V⊆U: each ρ ∈RpreservesV-validity}.

Note that while the local range of a rule or set of rules is a set of
valuations, the global range is a set of sets of valuations. But in
some cases there is an especially simple relationship between the two,
namely that *Glo*(*R*) = ℘(*Loc*(*R*)). This
happens whenever *R* comprises only zero-premiss rules, and, more
interestingly, when the rules in *R* are interderivable with a set of
such rules, given the structural rules (Id), (Weakening), and
(Cut),^{[13]} as in
the case of the natural deduction rules for conjunction *R* = {(∧I), (∧E)} or the corresponding sequent calculus rules {∧ Right, ∧ Left} where *Loc*(*R*) comprises precisely the
∧-boolean valuations. Evidently, the strong classicality
phenomenon of Section 2 is making
its presence felt here, and we should contrast the situation with
disjunction
(in Set-Fmla). For
*R* = {(∨I), (∨E)}, *Loc*(*R*) is the class of
∨-boolean valuations, but the same is the case if
(∨E)—the rule of current interest since it has no
zero-premiss equivalent—is replaced by another which is far from
interderivable with it (even given the structural rules and (∨I)):

A≻BA∨B≻B

Thus the local range is not very informative as to the precise
significance of such a rule as (∨E) and it is natural if
one thinks of the rules as giving the meanings of the connectives they
govern (a perspective mentioned in Section 1) and wants
to distill this information into a semantics in terms of valuations,
to consider the global range. To do so we make use of the partial
order ≤ on the underlying set *U* of valuations (for the language
under consideration) defined by: *u* ≤ *v* if and only for all
formulas φ, if *u*(φ) = T then *v*(φ) = T.^{[14]} As Garson (2001) observes, the global range of
the introduction and elimination rules for ∨ (and again we could
equally well cite the sequent calculus rules in this connection) is
the set of all *V* ⊆ *U* such that for all *u* ∈ *V* and all
formulas φ and ψ, the following condition is
satisfied:^{[15]}

[∨] _{Garson}u(φ ∨ ψ) = T iff for all formulas χ such thatu(χ) = F, there existsv∈Vwithu≤v,v(χ) = F and eitherv(φ) = T orv(ψ) = T.

While we cannot go into the details of Garson's program here, its aim
should be noted: to arrive at the ‘natural semantics’
implicit in the rules governing a connective # (here assumed *n*-ary)
in the sense of a characterization of their global range as comprising
those *V* of valuations satisfying a condition of the form:
∀*u* ∈ *V*, *u*(#(φ_{1},…,φ_{n})) = T if and only if
…, where the dots are replaced by something which would render
this biconditional suitable as a clause in the inductive definition of
truth on a valuation *u* ∈ *V*. Garson worries about the use of
< on this score in what we are calling
[∨]_{Garson} (and has various
suggestions to offer), as well as about the explicit quantification
over χ which introduces an potential element of circularity or
impredicativity (in particular because χ need not have a lower
complexity—be constructed with the aid of fewer primitive
connectives, that is—than φ ∨ ψ itself). One way of
getting rid of the latter feature is to massage
[∨]_{Garson} into the
following equivalent form, which makes extensive use of conjunctive
combinations of valuations, and uses the notation |χ| which, for
a formula χ, denotes the set of *v*∈ *V*
such that *v*(χ) = T:

u(φ ∨ ψ) = T iff ∃U_{0},U_{1}⊆V:U_{0}⊆ |φ|,U_{1}⊆ |ψ|, andu= ∏U_{0}⋅ ∏U_{1}.

In place of this last “∏ *U*_{0} ⋅ ∏ *U*_{1}”, we
could equally well have written “∏ (*U*_{0} ∪
*U*_{1})”, of course. The form as written is more suggestive of the
many variations on the semantic treatment of disjunction in the
literature—the Beth semantics for intuitionistic logic, for
instance, and the treatment of disjunction for relevant and other
substructural logics to be found in Ono and Komori (1985),
Orłowska (1985, 455), Bell (1986), and Humberstone (1988), to cite only
some papers from the 1980s. We will not discuss other connectives from
this perspective here, or comment further on Garson's program. See
Garson (2001) for a discussion of the global ranges of
the usual sets of introduction and elimination rules for →, ¬
and ↔.^{[16]}

The idea which, has just re-surfaced in our discussion, apropos of
Garson, that the natural deduction introduction and elimination rules
together, or perhaps just the introduction rules alone
(Gentzen's preferred position),^{[17]} somehow fix the meanings of the
connectives they govern, often goes hand in hand with considerations
of the existence and the uniqueness of a connective answering to the
said rules, considerations first raised by H. Hiż, according to
Belnap (1962). Belnap was replying to Prior
(1960) in which the idea that any old set of rules
could be stipulated to fix the meaning of a proposed connective, was
contested with the aid of the subsequently much-discussed
example Tonk, a supposed binary
connective for which an introduction rule allowed the passage from
(arbitrary) φ to
φ Tonk ψ, as with
∨-introduction from the first disjunct, and an elimination rule
allowed the passage from
φ Tonk ψ to ψ, as
with ∧-elimination to the second conjunct. Since this allowed,
by successively applying (Tonk-I)
and then (Tonk-E), any ψ to be
derived from any φ, Prior took this to show that no coherent
meaning had been conferred on Tonk
by the proposed rules, and to refute the general thesis that the
meaning of a connective can be fixed by means of rules governing
it. (Prior saw Popper and Kneale as the targets of his argument, which
does not touch Gentzen's view that the meaning is fixed by the
totality of introduction rules for a connective, coupled with the fact
that these are the totality of such rules. Since there is only the one
way of introducing Tonk, namely
that provided by (Tonk-I), where the
premiss is φ the meaning of
φ Tonk ψ is simply that
of φ and the *appropriate* elimination rule is not that
given by Prior but rather one licensing an elimination to the
*first* “tonkjunct”.) Belnap responded that what was wrong
with Prior's Tonk rules was that a proof system antecedently
endorsed for the connectives in use would be non-conservatively
extended by the enrichment of the language with Tonk in tandem
with supplementing the proof system with the (Tonk-I) and then
(Tonk-E), since they allow the derivation of the sequent φ
≻ ψ even for Tonk-free φ,
ψ.^{[18]} Later in this section, we
return to Tonk and examine its behaviour in a sequent calculus
setting.

Of course, one expects the addition of new rules to make for the
provability of new sequents, but the extension should be
*conservative* in the sense that no new sequents exclusively
involving the *old* vocabulary become provable. Belnap's
discussion suggests that the existence of a connective obeying certain
rules should be conceded by someone endorsing a particular proof
system if and only if the addition of that connective to the language
and those rules to the proof system yields a conservative
extension. (There are other things one might mean by the question of
whether from the perspective of a given logic, a connective with
such-and-such logical powers exists—most evidently, whether or
not such a connective is already definable in the logic. Showing that
rules ascribing the powers in question would extend the logic
non-conservatively would be one dramatic way of showing no such
connective to be definable, but since it could fail to be definable
while still being conservatively addable, the questions need to be
kept distinct.)

One snag for such a suggestion is that examples can be found (see
Section 6 of Humberstone (2001)) where each of two connectives with
associated sets of rules can be such as to yield a conservative
extension of a proof system, apparently requiring the answer
“yes, there does exist such a connective” in each case,
while extending the proof system with both at once is
non-conservative—undermining this double endorsement. We may
conclude that while yielding a non-conservative extension of a
favoured proof system is sufficient to denying the existence or
intelligibility of a connective with such-and-such logical properties,
there may be other grounds for returning such a verdict. Boolean
negation comes under just such a critique in Chapter 5 of Priest
(2006), where
such problems for Belnap's account are raised
and references to the technical literature establishing the
conservativity of extending certain logics in the relevance tradition
with suitable principles governing this connection can also be
found. (For an earlier presentation of this point, see also Belnap and
Dunn (1981), p. 344, and for a related discussion, Restall
(1993), esp. §5.) It should be noted also that
the conservativity test raises issues of its own, independently of the
question of whether passing it is necessary and sufficient (and what
we have been querying is the sufficiency half of this package) for
acknowledging the existence of a connective. For example, there are
sentential connectives which conservatively extend intuitionistic
*propositional* logic in any of its conventional presentations
as a proof system, but non-conservatively extend intuitionistic
*predicate* logic. The best known is dual intuitionistic
implication (see López-Escobar
(1985)).^{[19]} Returning to the
conservative extension criterion of existence, let us note that a
related connective, dual intuitionistic negation, can be added
conservatively to intuitionistic predicate (and *a fortiori*
propositional logic) but already at the propositional level produces
violations of the *disjunction property*—that
≻ φ ∨ ψ should only be provable if at least one
of ≻ φ, ≻ ψ, is provable—which on
some accounts (e.g., Gabbay (1977) and (1981)) renders it illegitimate
from an intuitionistic perspective. Finally, it should be observed
that an adherent of intuitionistic logic might be expected to take a
dim view of a proposed new connective with rules conservatively
extending that logic but not conserving synonymies, in the sense that
formulas in the original language which are synonymous according to
⊢_{I L}
are no longer so according to the
consequence relation associated with the extended proof system. This
happens in at least the most straightforward extension of
intuitionistic logic by what is called
*strong negation*.^{[20]}

The issue of uniqueness, as opposed to existence, of connectives (in
the logically loaded sense of the term, naturally) is best approached
by considering when a collection of rules governing a connective
*uniquely characterize* that connective in the sense that if
accompanied by corresponding rules governing a ‘duplicate’ version of
that connective (of the same arity) result in the synonymy of
compounds formed from the appropriately many components and the
original connective with compounds formed from the duplicate
connective from the same components. In a congruential setting the
reference to synonymy can be replaced by one of equivalence, in the
sense of ⊣⊢-equivalence in the consequence relation or
Set-Fmla framework, with suitable variations for other
frameworks. Without the assumption of congruentiality one should
perhaps distinguish unique characterization to within synonymy from
unique characterization to within equivalence. And because the rules
in question may not be pure, but involve other connectives, it may be
necessary to speak of # being uniquely characterized (by the given
rules) *in terms of* the other connectives involved, so that
when the reduplicated system is considered, # is replaced by its
duplicate, #′, say, while the other connectives remain as they
are. We will not be considering these complications here. It is easy
to see that the natural deduction rules for the various connectives
given in Section 2 as well as the sequent calculus
rules all uniquely characterize the connectives they govern. (Note
that, since this includes the introduction and elimination rules for
→, which are shared by any natural deduction system for
intuitionistic logic, the rules may uniquely characterize a connective
which is not fully determined—not amenable to a truth-functional
interpretation. The same applies in respect of negation, though we did
not cover this in our discussion of natural deduction.) This feature
is sometimes regarded as necessary for the connectives concerned to be
regarded as logical constants, although, as explained in Section
1, we are not pursuing that issue
here.

A related issue involving both uniqueness and existence does deserve
attention, however, arising over competition between logics, differing
on the rules governing a connective where the first logic has rules
which suffice to characterize uniquely a connective and the second has
all these and more governing that connective. The standard example of
this is provided by intuitionistic logic and classical logic in
respect of negation. The fact that the classical rules go beyond the
intuitionistic rules and yet the latter already suffice for unique
characterization means that there is no ‘live and let
live’ option available for the intuitionist, of recognising that
the intelligibility of two distinct negation connectives, classical
and intuitionistic, which can cohabit in the same proof system each
with their respect rules. This should be possible according to those
who think that any difference in the logical principles governing a
connective reflect a difference as to what is meant by the connective
(W. V. Quine's view—see Section 5
below), since once a difference in meaning is acknowledged, the
previously ambiguous connective, “¬”, in the present
instance, should be replaceable by connectives registering the
disambiguation—¬_{i} and
¬_{c}, say—and the
intuitionistic and classical logician should be able to use whichever
they take a special interest in, with no further fear of talking at
cross-purposes in each other's company. But of course this is not a
possibility, since the rules governing
¬_{i} characterize it uniquely
and those governing ¬_{c}
duplicate them—think of
¬_{c} as
¬_{i}′, in terms of the above
discussion—*and do more* (securing the provability of
¬_{c}¬_{c}*p*
≻ *p*, for example). The result is that the two connectives form
equivalent compounds in the reduplicated system and all of the
intuitionistically unwanted properties of ¬_{c} spill over to
infect ¬_{i}: we can now prove
¬_{i}
¬_{i}*p* ≻ *p*, and so
on. The conciliatory offer of a purportedly disambiguating notation
along these lines is one no advocate of intuitionistic logic can
afford to accept. (By imposing restrictions on some of the
intuitionistic rules, one can produce candidates for the title of
‘combined classical-and-intuitionistic proof
system’—as, for example, in Fariñas del Cerro and
Herzig (2001)—but, especially on views
according to which the rules of one's logic articulate the meanings of
the connectives they govern, resistance to any such restrictions is
entirely understandable. Imagine an encounter with aliens in which
they told us that to explain what they meant by some items of their
vocabulary we had to be prepared to restrict the principle that
ψ followed from φ* and *ψ for certain choices
of ψ involving translations of that vocabulary into
English.)

The above line of thought points in some interesting further
directions. For example, the intuitionistic natural deduction rules
governing ∧
(in Set-Fmla),
which are the same as the classical ones, are themselves stronger than
the following rules:^{[21]}

χ ≻ φ χ ≻ ψ χ ≻ φ ∧ ψ

χ ≻ φ ∧ ψ χ ≻ φ

χ ≻ φ ∧ ψ χ ≻ ψ

We can take the intuitionistic (and classical) rules to differ from
these in replacing the formula variable χ with a set-of-formulas
variable Γ. The latter set is stronger in the sense that there
are sequents provable with its rules (e.g., *p*, *q* ≻ *p* ∧ *q*)
which are not provable on the basis of the displayed rules (together
with the standard structural rules, if desired). But clearly the
displayed rules uniquely characterize ∧, since, once we
reduplicate them for ∧′, we can take χ in the second and
third (collectively, (∧E)-style) rules to be φ ∧
ψ, to derive both φ ∧ ψ ≻ φ and φ
∧ ψ ≻ ψ, from which by the reduplicated version of
the first (“(∧′I)”) rule we derive φ ∧
ψ ≻ φ ∧′ ψ—with a similar derivation for
the converse sequent.

This example may seem far-fetched from the point of view of
motivation, but the dialectical upshot is clear—if the ¬ rules
are objectionable in being stronger than needed for unique
characterization, the same holds for the intuitionistic ∧ rules;
in any case there are in fact variations on this theme is that have
been seriously proposed as weakenings of classical (or intuitionistic)
logic in the construction of quantum (or ortho-) logic, in which the
rule (∨E) is subjected to restrictions blocking arbitrary sets of
side-formulas—the Δ and θ of our formulation of this
rule in Section 2—exactly as with the ∧ rules
inset above (which would be fully general in the respect at issue if
our chosen framework were what would naturally be called Fmla-Fmla, a framework occasionally used for logical
investigations).^{[22]} One could explore the idea of a requirement that all
rules should be maximally general in respect of side formulas, to
resist such pressures to weaken the intuitionistic rules, though what
this amounts to depends on the logical framework, and it may be hard
to hold the line without succumbing to a similar pull in the direction
of generality which would take us over from Set-Fmla to
Set-Set.

The sensitivity to logical framework of issues of unique characterization is well illustrated with the case of modal logic. It is possible to fabricate special purpose logical frameworks which allow for the formulation of rules uniquely characterizing ☐ in all normal modal logics (Blamey and Humberstone 1991), but without such exotic devices, unique characterization is limited to the case in which ☐ is fully determined—amongst the normal modal logics, then, just to that for which ⊢ ☐ φ for all φ, and that in which φ ⊣⊢ ☐ φ for all φ. (Note 31 is relevant here. The “⊢” just used stands for the local consequence relation concerned so we could equally well have written, “⊢ φ ↔ ☐ φ”.)

This same framework-sensitivity issue can be illustrated by
considering substructural logics in multiset-based
frameworks. Classical Linear Logic has two versions of the
propositional constants (or 0-ary connectives) ⊥ and ⊤,
called (in the relevance tradition—see
Avron (1988)—rather than in the notation of the seminal paper Girard (1987)) *F* and *f* in the former case
and *T* and *t* in the latter, with the
following Mset-Mset
rules, the Greek letters now standing for finite multisets of formulas
(and commas for multiset union):

( FLeft)Γ, F≻ Δ( TRight)Γ ≻ T, Δ( tLeft)

Γ ≻ Δ Γ, t≻ Δ( tRight)≻ t( fLeft)f≻( fRight)

Γ ≻ Δ Γ ≻ f, Δ

If we were doing classical logic with multisets, we would need the
structural rules (Id), (Weakening) and (Cut) from Section
2, set-variables now reinterpreted as multiset
variables, and the further rule of Contraction to allow double
occurrences of formulas on the left or right of the ≻ to be
replaced by single occurrences. For linear logic we retain only (Id)
and (Cut), disallowing weakenings and contractions, which has the
effect that there are now also two versions of conjunction and of
disjunction, multiplicative and additive—into which distinction we
cannot go further here, except to say that for binary connectives it
lines up with the intensional/extensional distinction in the relevance
tradition.^{[23]}
A sequent calculus presentation of the favoured logic from this
tradition, **R**, can be obtained from that for Classical Linear
Logic (without two of its characteristic 1-ary connectives, called
exponentials, which are there to mark formulas for which weakening and
contraction are permitted) by allowing contraction (while still
banning weakening).^{[24]} We return to **R** in Example 6 of
Section 4; it is for the sake of that return that the
rules governing *t* and *T* have been given here alongside those for
*f* and *F*, the latter pair being of current interest. Notice that
even though there is only one rule for *F*, namely a left insertion
rule (with no premiss sequents), this rule suffices to characterize
*F* uniquely, since in the combined language with *F*′ governed by a
like rule we can prove—just by appealing to the appropriate
instance of the rule *F* ≻ *F*′ and *F*′ ≻ *F*. The
weaker system of Intuitionistic Linear Logic is cast in the framework
Mset-Fmla_{0} disallowing the appearance of more than one
formula on the right, and this application of the left rule for *F* is
available there, as also in Mset-Fmla, requiring exactly
one formula on the right (since we have exactly one formula—*F* or
*F*′—there in the present case). When we turn to *f*, again we find
that the rules uniquely characterize this connective, since (*f* Left)
provides the premiss for an application of (*f*′ Right) with Γ
consisting of (one occurrence of) *f* and Δ empty, yielding *f*
≻ *f*′, with the converse proved similarly. (Note that while
*F* ≻ *f* is a special case of the (left) rule for *F*, there
is no proof of the converse sequent. One cannot begin with *f*
≻ and then weaken in an *F* on the right, since we do not
have Weakening.^{[25]})
Again, these derivations go through in Mset-Fmla_{0}, but this time we note that they do not go through in Mset-Fmla, since the left insertion rule has no applications in
that framework. When attention is restricted to Mset-Fmla,
there is no set of pure rules available which uniquely characterize
*f*. The Mset-Mset (or Mset-Fmla_{0}) sequent
calculus rules for ¬, which can be taken to be those given in
Section 2 with the set-variables taken as
multiset-variables (suitably restricted for Mset-Fmla_{0}), uniquely characterize ¬, but have no direct
application in Mset-Fmla or any other framework with a
fixed multiset cardinality on the right (or the left, for that
matter). We can say in this framework that the rules governing ¬
uniquely characterize this connective *in terms of* → and
*f*—on the basis of which it would indeed be explicitly definable
(¬ φ as φ → *f*), but since here *f* is not uniquely
characterized (even though → is), this still leaves ¬ without
its (perhaps) expected uniqueness, as is even more evident in Fmla. These considerations continue to apply in the presence of
contraction (and ∧/∨) distribution, accounting for the
complaint sometimes made that relevant logic does not operate with a
uniquely characterized negation. The point here has not been either to
press or to reply to such complaints, but to note their framework
sensitivity: they do not arise for Mset-Mset or Mset-Fmla_{0}.

The existence and uniqueness questions for connectives turn out in the
case of sequent calculus rules to be intimately connected with the
structural rules (Id) and (Cut). Let us call the left and right
insertion rules governing an *n*-ary connective #—or derivatively,
the connective thereby governed—*Id-inductive* if the
provability of #(φ_{1},…,φ_{n}) ≻
#(φ_{1},…,φ_{n}) follows from the provability of each φ_{i}
≻ φ_{i} (*i* = 1,…,*n*) by means of those rules (without
assistance from other structural rules). Thus means that no provable
sequents are lost if the rule (Id) itself is restricted to having
applications φ ≻ φ for φ a propositional
variable. And let us say that the rules in question are
*Cut-inductive* if an application of the rule (Cut) with the
formula #(φ_{1},…,φ_{n}) as cut-formula (i.e., as the φ
of our representation of the rule (Cut) in Section 2 or
the analogous rule in Set-Set or the multiset based
frameworks) can be dispensed with if the premiss sequents for the
application of the rule have been derived by (# Left) and (# Right)
and (Cut) is allowed with the components φ_{1},…,φ_{n} as
cut-formulas (but again, not allowing the use of other structural
rules).

All the connectives of the standard sequent calculi, including the
cases of classical and intuitionistic linear logic, are inserted by
rules which are both Id-inductive and Cut-inductive, as we illustrate
here with the rules for →. For definiteness, think of these as
the Left and Right rules of classical linear logic. To show
Id-inductivity, we must show that
φ → ψ ≻ φ → ψ
is provable using the → rules, given (‘the
inductive hypothesis’) that
φ ≻ φ and
ψ ≻ ψ are.
(We write φ and ψ rather than
φ_{1} and φ_{2}.) From the last two sequents, (→ Left)
delivers: φ → ψ, φ ≻ ψ. From this an
application of (→ Right) gives the desired
conclusion:
φ → ψ ≻ φ → ψ.

For Cut-inductivity suppose (1) Γ, Γ′, φ → ψ ≻ Δ, Δ′ has arisen by (→ Left) from (1a) Γ ≻ φ, Δ and (2a) Γ′, ψ ≻ Δ′, while (2) Γ′′ ≻ φ → ψ, Δ′′ has arisen by (→ Right) from (2a) Γ′′, φ ≻ ψ, Δ′′. We need to show that a cut on (1) and (2) with cut-formula φ → ψ can be avoided on the hypothesis that we can help ourselves to cuts with cut-formulas φ, ψ. From (1a) and (2a), cutting on φ, we get (3) Γ, Γ′′ ≻ ψ, Δ′′. Then cutting on ψ, (2b) and (3) deliver the desired conclusion: Γ, Γ′, Γ′′ ≻ Δ, Δ′, Δ′′, without the need to use φ → ψ as a cut-formula with (1) and (2) as premisses.

Cut-inductivity is one aspect of a procedure which has traditionally
been a professional obsession of proof-theorists: the provision of
‘Cut Elimination’ theorems to the effect that this or that sequent
calculus with the rule (Cut) can have this rule removed without the
loss of any provable sequents. Obviously more is needed to show this
than the Cut-inductivity of the (pairs of) rules, since there are
also, *inter alia*, potential interactions with the other
structural (and operational) rules to be taken care of. Such results
are of interest because much follows from the subformula property—any sequent provable has a proof in which the only formulas to appear
are subformulas of those in the sequent concerned—which itself
follows from the result since all the remaining rules insert material
rather than remove it in the course of the proof.^{[26]}
But from the perspective of the philosophy of logic, Cut-inductivity
and its somewhat less well-publicized companion, Id-inductivity,
should be considered on a par, as the following somewhat
impressionistic semantical remark (Girard, Taylor and Lafont (1989), p. 31) of
Girard suggests, apropos of (Cut), in which *C*
represents the cut formula (and (Id) with *C* on the left and right of
the “≻”):

The identity axiom says thatC(on the left) is stronger thanC(on the right); this rule [ = (Cut)] states the converse truth, i.e.C(on the right) is stronger thanC(on the left).

The point of isolating Id- and Cut-inductivity here, however, has less to do with matters of truth (on which, and for further references to Girard on which, see further Hösli and Jäger (1994)) than with uniqueness and existence—not that the semantical issues are unconnected with this, of course. The former link is perhaps already clear. At the last step of establishing Id-inductivity for →, above, we passed from φ → ψ, φ ≻ ψ, obtained by (→ Left), to φ → ψ ≻ φ → ψ, by (→ Left): but this is just what the uniqueness proof in the reduplicated sequent calculus would do, except that the final step would be an application of (→′ Right) and would establish φ → ψ ≻ φ →′ ψ, with the converse proved by interchanging the rules used for → and →′.

To see the connection between existence—or more specifically, the conservative extension side of the existence question—convert the natural deduction rules for Tonk given earlier into sequent calculus rules (here for Set-Fmla):

(Tonk Left)

Γ,ψ ≻ χ Γ, φ Tonk ψ ≻ χ

(Tonk Right)

Γ ≻ φ Γ ≻ φ Tonk ψ

By applying the right insertion rule with premiss sequent (courtesy of (Id)) φ ≻ φ and the left rule with premiss ψ ≻ ψ we arrive at the point at which our natural deduction derivation started to go bad: we have:

φ ≻ φ Tonk ψ and φ Tonk ψ ≻ ψ

and an appeal to (Cut) here, with
φ Tonk ψ as cut-formula,
amounts to bringing home the bad news.^{[27]}
But of course this example simply makes
vivid the fact that the Tonk rules
in play are not Cut-inductive and there is no Cut-free proof of the
desired sequent (at least in general: take φ = *p*, ψ =
*q*). This is hardly surprising since the subformula property which Cut
Elimination secures itself guarantees that the rules for one
connective conservatively extend the subsystem comprising the rules
for the remaining connectives. In the present case, the point of the
example was that there was no proof of, for example, *p* ≻ *q*
which did not detour—in violation of the subformula
property—via formulas
involving Tonk.

We can look at what is roughly speaking a mirror-image of the Tonk example to reinforce the contrast between Cut-inductivity and
Id-inductivity and their respective links with existence and
uniqueness, and then sort out the “roughly speaking” part of this
description, which will show them as providing, rather, special—one
might say canonical—cases of the satisfaction of the existence ( =
conservative extension) and uniqueness conditions. We conduct the
discussion in Set-Set rather than Set-Fmla to
avoid a mass of complications. The above Tonk rules combine a
left insertion rule familiar from the sequent calculus treatment of
∧ and a right insertion rule familiar from that of ∨; their
reformulation for the change of framework just announced requires an
idle Δ parameter on the right of the sequents in the premisses
and conclusions of the rules as schematically presented above
(replacing the χ in the one case, accompanying the displayed
succedent formulas in the other). For our mirror image
example,^{[28]} consider a binary
connective Knot with rules combining instead the left insertion
rule for ∨ with the right insertion rule for ∧:

(Knot Left)

Γ,φ ≻ Δ Γ,ψ ≻ Δ Γ, φ Knot ψ ≻ Δ

(Knot Right)

Γ ≻ φ, Δ Γ ≻ ψ, Δ Γ ≻ φ Knot ψ, Δ

These rules are Cut-inductive, as the interested reader may confirm,
but not Id-inductive, since they do not uniquely characterize Knot. To see this, note that both are derivable rules for ∧ and
also for ∨, and of course we cannot prove φ ∨
ψ ≻ φ ∧ ψ in general. By contrast, neither
∧ nor ∨ satisfies *both* of the left and right rules
for Tonk. The upshot of the Knot rules is best visualised
in terms of the zero-premiss rules (sequent-schemata) with which they
are interderivable given the structural rules:
φ, ψ ≻ φ Knot ψ (the right rule) and
φ Knot ψ ≻ φ, ψ (the left rule). An
incidental point: since these are the sequent formulations of
conditions induced by the determinants ⟨*T*, *T*, *T*⟩ and ⟨*F*, *F*, *F*⟩ one might
think of them together as expressing idempotence and therefore as
being interchangeable with the simpler pair φ
≻ φ Knot φ and
φ Knot φ ≻ φ. But while the latter two can
be derived from the earlier pair (just take ψ as φ), the
converse derivation is not possible without further principles—conspicuously, a sequent version of the extensionality conditions on
generalized consequence relations mentioned in Section
2. We could call binary # for which φ, ψ
⊩φ # ψ and conversely, *strongly idempotent*
according to the generalized consequence relation ⊩. Then
∧ and ∨ are not the only boolean candidates for a
connective behaving as the Knot rules dictate—strongly
idempotent according to the generalized consequence relation ⊩_{CL}
determined by the class of boolean valuations: we also have the binary
projection to the first coordinate (φ # ψ equivalent to
φ), to the second coordinate (φ # ψ equivalent to
ψ), and venturing further afield, the Set-Fmla
hybrids of any two of these connectives. Of these perhaps the most
interesting is the hybrid of the two projection connectives, since
this—let us write is as “•”—can be used to define all
other hybrids explicitly from the connectives to be hybridized; for
example if we want the common properties (as explained in Section
2) of → and ∨, we have them in the following
compound of
φ and ψ: (φ → ψ) • (φ ∨ ψ).^{[29]}

If the treatment of Knot were genuinely the mirror image of that of
Tonk, we would be able to say that just as the Knot rules
are Cut-inductive but not Id-inductive, so the Tonk rules are
Id-inductive but not Cut-inductive. (The twofold demand of Cut- and
Id-inductivity is a promising candidate for the notion of
*harmony* as it applies to the relation between right and left
insertion rules in a sequent calculus.^{[30]}) And while the failure of
Cut-inductivity was indeed noted, nothing was said on the subject of
Id-inductivity for these rules. Spoiling the intended duality, the
situation is that the Tonk rules fail on this score too—though this can be remedied by a simple variation. First, we should
note that Prior's point could equally well have been made by combining
the ∨ introduction rule from the *second* (rather than the
to first) disjunct with the ∧ elimination rule to the
*first* (rather than to the second) conjunct. Call the
connective with these
rules Tonk*

Tonk* Left

Γ,ψ ≻ Δ Γ, φ Tonk* ψ ≻ Δ

Tonk* Right

Γ ≻ ψ, Δ Γ ≻ φ Tonk* ψ, Δ

This suffers from the same defects in respect of the analogy with Knot, so, finally, let Tonk^{+} be governed by the rules for
Tonk and also the rules for Tonk*. It is Tonk^{+}
that has Knot as its mirror image, since just as we Knot
has the full force of the (∧ Right) and (∨ Left) rules,
Tonk^{+} enjoys the full force of the (∧ Left) and (∨
Right) rules. (So here we arrive at the reason that Knot was
originally described as only *roughly speaking* a mirror image
of Tonk.) And this time the rules are indeed Id-inductive
without being Cut-inductive. Indeed, Id-inductivity is secured by two
independent routes: by the interaction of the Tonk Left rule for
Tonk^{+} with the Tonk* Right rule to Tonk^{+}, and
then again by the interaction of Tonk^{+}‘s Tonk* Left
rule with its Tonk Right rule. (To give the former in more
detail, by way of example, the Tonk Left rule for Tonk^{+}
gives φ Tonk^{+} ψ ≻ψ from (Id)-for-ψ,
from which applying the Tonk* Right rule for Tonk^{+} we
obtain φ Tonk^{+} ψ ≻ φ Tonk^{+} ψ.)

Let us sum up the lessons for the relations between existence and Cut-inductivity and between uniqueness and Id-inductivity that can be extracted from the foregoing discussion. Sequent calculus rules which are Id-inductive uniquely characterize the connective they govern, but not necessarily conversely: the Tonk rules are not Id-inductive, but they do uniquely characterize Tonk, since they are already so strong that any sequent with at least one formula on the left and at least one on the right can be proved with their aid, and hence, in the reduplicated system with Tonk′, we can certainly prove φ Tonk ψ ≻ φ Tonk′ ψ and the converse sequent—without even using the Tonk′ rules! (This is the kind of thing we have in mind in calling this a non-canonical form of uniqueness.) On the ‘existence’ side: it is not clear to the author whether any Cut-inductive rules governing a connective guarantee that they are conservative over the (system consisting of the) rules governing other connective—since Cut Elimination may fail for other reasons; nor do the examples just considered settle the converse, as we see from the following tabular summary of our findings, in which (a “Yes” by) “unique” means the rules uniquely characterize the connective, and “conservative” means that we have a conservative extension of any system of sequent-calculus rules not involving the given connective; Tonk* is not listed separately since it patterns the same way as Tonk:

Id-inductive Unique Cut-inductive Conservative Tonk No Yes No No Tonk ^{+}Yes Yes No No Knot No No Yes Yes

That is enough on the subject of Tonk and variations. We will
make one point about unique characterization and then give some
examples of non-conservative extension which will involve introducing
the important concepts of *special* and *universally
representative connectives*. Some further ‘existence’ examples are
deferred to the following section.

The above discussion has included such phrases as “the hybrid of
∧ and ∨”; the totality of principles governing such a
connective being those common to ∧ and ∨, we obviously do
not have a connective uniquely characterized by those principles—in
which case is it legitimate to speak of *the* hybrid…? I
take this way of speaking to be as harmless as speaking of “the
☐ operator of **S5**”, where again the logical principles
involved are insufficient to characterize the connective uniquely—in terms of the boolean connective or otherwise—at least if we are
Fmla, Set-Fmla
or Set-Set,
and in the latter cases, working with the so-called *local*
(generalized) consequence relation.^{[31]} (See
Section 5 for the distinction between
the local and the global consequence relations determined by a class
of models.) One could say that in the monomodal, unreduplicated case,
we have a single connective playing a certain role and we can refer to
it in that context as *the* connective playing the role in
question, while conceding, as we learn from the reduplicated case,
that this is a role that can be played simultaneously (i.e. in a
single logic) by several non-equivalent connectives. This is the kind
of situation Łukasiewicz had in mind when, in
his (1953), he likened certain pairs of connectives to
“twins who cannot be distinguished when met separately, but are
instantly recognised as two when seen together”.

We saw earlier that combining the intuitionistic and classical rules
in a single proof system did not work well, and another suggestion
that comes to mind is that we could perhaps extend intuitionistic
logic with a new connective, 1-ary #, say, which does what according
to the intuitionist, negation (¬) itself cannot do: serve as a
left inverse for ¬ in the sense of rendering equivalent
#¬φ and φ for all φ.^{[32]}
This again represents
an offer the intuitionist cannot afford to accept, assuming that any
(the rules for any) proposed new connectives should conserve synonymy,
as suggested above. (The even stronger demand that the consequence
relation associated with the extended proof systems should be
congruential does not seem unreasonable.) The reason is that ¬*p*
and ¬¬¬*p* are intuitionistically equivalent (and thus
synonymous according to ⊢_{I L}) so the conservation of synonymy
requires that #¬*p* and #¬¬¬*p* should be
equivalent—since we have this ‘Law of Triple Negation’ equivalence
in intuitionistic logic—from which the left inverse condition above
delivers, taking φ as ¬*p* in the former case and as
¬¬¬*p* in the latter, the equivalence of *p* with ¬¬*p*—a straightforwardly non-conservative extension, then. (A similar
argument shows that adding binary # with φ # (φ
↔ ψ) equivalent to ψ—a possibility
suggesting itself since, by contrast with the classical case,
↔ itself cannot serve as # here—is also
non-conservative.) Nor, incidentally, could the intuitionist
acknowledge the existence of a *right* inverse for negation, a
# with arbitrary φ equivalent to ¬ # φ, as this makes
every formula equivalent to a negated formula and hence to its own
double negation. This argument, which is well known (see, e.g., Koslow
(1992), p. 97), does not even require the
‘conservation of synonymies’ constraint.

These last non-existence proofs force our attention on the fact that
negated formulas are ‘special’ in intuitionistic logic in that each
such formula is equivalent to its own double negation while this is
not so for arbitrary formulas. Let us back up and give a clear
explanation of the notion of specialness as well as a contrasting
notion, in slightly more general terms; for definiteness only, we
tailor the account to Set-Fmla. We also begin with the
contrasting notion. Say that an *n*-ary relation *R* amongst the
formulas of a language is *universally representative*
according to a consequence relation on that language if for any
formulas φ_{1},…,φ_{n} of the language, there are formulas
ψ_{1},…,ψ_{n}, respectively synonymous with
φ_{1},…,φ_{n} according to ⊢, with
*R*(ψ_{1},…,ψ_{n}). (If *n* = 1, we speak instead of a
universally representative *class* of formulas—a class of
formulas such that every formula is synonymous with some formula in
the class; likewise in the case of the notion *special*,
defined below.) If ⊢ is congruential then we can replace
“synonymous” in this definition with “equivalent” (i.e. such that
φ_{i} ⊣⊢ ψ_{i}, for *i* = 1,…,*n*). Say that *R* is
*special* according to ⊢ if there is some Set-Fmla sequent
σ(*p*_{1},…, *p*_{n})
constructed using not
necessarily just the variables exhibited, for which for any formulas
ψ_{1},…,ψ_{n} with *R*(ψ_{1},…,ψ_{n}), we have
σ(ψ_{1},…,ψ_{n}) ∈ ⊢,
while we do not have, for
all formulas
φ_{1},…,φ_{n},
σ(φ_{1},…,φ_{n}) ∈ ⊢.
For substitution-invariant ⊢, this last condition
can be simplified to: σ(*p*_{1},…, *p*_{n}) ∉
⊢.^{[33]}
A consequence
of these definitions is that no relation is both universally
representative and special according to any given ⊢. A simple
example of a special relation—according to more or less any
‘naturally occurring’ consequence relation with ∨ in its language
is relation *R* obtaining between φ and ψ just in case
ψ is a disjunction of which φ is a disjunct. Here take
σ(*p*, *q*) as *p* ≻ *q* which is (e.g., in intuitionistic
logic, though classical logic or any of many others would do as well)
unprovable—i.e., not *p* ⊢_{I L} *q*—while φ ⊢_{I L}
ψ for all φ, ψ for which *R*(φ, ψ). A more
interesting example is the relation of being (intuitionistically)
*head-linked*, i.e., that relation *R* such that for all
φ, ψ, *R*(φ, ψ) if and only if:

for some formulas χ, φ_{0}, ψ_{0}, we have φ ⊣⊢_{I L}φ_{0}→ χ and ψ ⊣⊢_{I L}ψ_{0}→ χ.

The analogous relation for classical logic would be universally
representative (and thus not special), at least in any fragment which,
as well as having →, has either ⊥ or ∧, since any
φ and ψ are here respectively equivalent to (φ → ⊥)
→ ⊥ and (ψ → ⊥) → ⊥, where they are linked
explicitly by having a common χ at the head of the main
implication. (Alternatively with ∧, replace ⊥ here by φ
∧ ψ.) But in the intuitionistic case the relation of being
head-linked is special since, for example, with Peirce's law
σ(*p*,*q*) = (*p* → *q*) → *p* ≻ *p*, we have σ(*p*, *q*)
intuitionistically unprovable while σ(φ,ψ) ∈
⊢_{I L} for any head-linked φ, ψ.

Concentrating now on 1-ary relations between—i.e., classes of—formulas, we transfer the above terminology to connectives by saying
that a connective is *universally representative* or
*special* according to a consequence relation if the class of
formulas with the given connective as main connective is itself a
universally representative class or a special class of formulas,
respectively, according to the consequence relation in question. While
every non-constant connective^{[34]}
is universally representative according to ⊢_{CL} our
discussion of inverses for intuitionistic ¬ shows that ¬ is
special according to ⊢, since we can take σ(*p*) as
¬¬*p* ≻ *p* and every substitution instance of this
IL-unprovable sequent in which *p* is replaced by a negated formula is
IL-provable. Negation is therefore not universally representative in
intuitionistic logic. (This means that we could usefully speak, from
the perspective of ⊢_{I L}, of negative *propositions*,
thinking of the equivalence class, with respect to the relation of synonymy, of
a formula as the proposition thereby expressed, and meaning by this
the equivalence classes of negated formulas. Classically, the
corresponding notion is not useful since all propositions would count
as negative, just as for both ⊢_{CL} and ⊢_{I L}, there
are no specifically conjunctive or disjunctive propositions.)

Although the properties of being special and being universally
representative according to a given logic are mutually exclusive, they
are not jointly exhaustive. For example, ⊥ is neither universally
representative nor special in minimal logic (according to the
consequence relation associated with the Johansson's
*Minimalkalkül*, that is). Nor does the connective
☐ have either of these properties in the modal logic **KD**
(sometimes just called **D**), though this example is far from
selected at random: in the sublogic **K** ☐ is special
(since ☐⊥ → φ is provable for all
☐-formulas φ, though not for arbitrary φ), and in
the extension **KD4** it is special again (since φ →
☐φ is provable for all ☐-formulas φ, though
not for all formulas).

## 4. Selected Existence Questions

### 4.1 Conditions not Corresponding to Rules (Example 1)

For conditions other than those corresponding to (closure under certain) rules we can also ask whether a connective satisfying the condition in question can safely—i.e. conservatively—be postulated to exist. One such condition is the following, here expressed as a condition on a consequence relation ⊢ with a binary connective # in its language, all variables understood as universally quantified:

Γ ⊢ φ # ψ if and only if Γ ⊢ φ or Γ ⊢ ψ.

It is because of the “or” on the right that this condition—more
precisely, its ‘only if’ half—does not correspond to a
sequent-to-sequent rule. Still we can ask whether some Set-Fmla proof system including the rule (∨E) (or the sequent
calculus rule (∨ Left)), e.g. for classical logic, has a
conservative extension to a proof system for a language with the
additional connective #, whose associated consequence relation
satisfies the above condition for this #.^{[35]}
To see that the answer to this
question is negative, note that the binary relation R_{⊢}
defined in terms of the #-supporting extension ⊢:
*R*_{⊢}(φ,ψ) iff φ ⊢ ψ, gives rise to a binary
relational connection (*L*, *L*, *R*_{⊢}) which provides conjunctive
(yes—*con*junctive) combinations on the left, since ⊢
is ∨-classical, φ ⋅_{L} ψ here being φ ∨
ψ, while the inset condition above (take Γ = {χ})
secures disjunctive objects on the right, with φ +_{R} ψ
being φ # ψ. (“*L*” and “*R*” subscripts from Section
1 restored here to prevent confusion, since the source
and target coincide.) Thus the crossover condition from Section
1 is satisfied, which says that if φ_{1} ⊢
ψ_{1} and φ_{2} ⊢ ψ_{2} then either φ_{1} ⊢ ψ_{2}
or φ_{2} ⊢ ψ_{1}. Taking φ_{1} = ψ_{1} = χ and φ_{2}
= ψ_{2} = χ′, we have the antecedent of this conditional (by
(Id)), so we conclude that for all formulas χ, χ′, either
χ ⊢ χ′ or χ′ ⊢ χ, which includes #-free
formulas of the original language (e.g., χ = *p*, χ′ = *q*) for
which neither is a consequence of the other by the original
consequence relation, showing any proposed extension along the above
lines to be non-conservative.

### 4.2 Pollard-Style Nonconservativity (Example 2)

This is really a whole range of examples. As we required (∨E) for
Example 1, here we require (→I). Let us call a consequence
relation ⊢ →-*introductive* if Γ, φ ⊢
ψ always implies Γ ⊢ φ → ψ.^{[36]}
Now suppose ⊢ ⊇
⊢_{I L} and ⊢ is →-introductive, and either (i) for
some *n*-ary connective #, ⊢ satisfies the conditions induced
(on consequence relations) by the determinant ⟨F,…, F,T⟩ (*n* occurrences of “F”) for
#, as well as some determinant with F as its final entry, or (ii)
⊢ satisfies the conditions induced by the determinant ⟨T,…,T, F⟩ (*n* occurrences of “T”) as well
as some determinant with T as final entry (for *n*-ary #). Then it
follows that ⊢ ⊇ ⊢_{CL}.

The implication from (i) to the conclusion that ⊢ ⊇
⊢_{CL} has been stressed in several publications by S. Pollard
(Pollard (2002), Pollard and Martin (1996)); there is a similar
proof for the implication from (ii). Neither proof is given here—only an illustration (of case (ii)). The truth-function known as
exclusive disjunction (or *xor*), which differs from that
associated on boolean valuations with ∨ (“inclusive
disjunction”) by having ⟨T, T, F⟩
instead of the determinant ⟨T, T, T⟩,
will serve as an example, since as well as the determinant just
mentioned, which has T's except for the final F, we have a T-final
determinant (e.g.) ⟨T, F, T⟩. The
induced conditions on generalized consequence relations for these two
determinants are respectively:

φ, ψ, φ # ψ ⊩ and φ ⊩ ψ, φ # ψ.

For the following, we will take it that, as well as →, the
connectives ∨, ¬, ⊤ and ⊥ are available,
though in fact the result just given requires is correct for any
fragment with → (with ⊢_{CL} understood as the
restriction of the classical consequence relation to the fragment in
question). Then we can write the induced conditions on our consequence
relation ⊢ as, respectively, rather than using a new schematic
letter on the right for the first, and converting the second into
conditional form:

φ, ψ, φ # ψ ⊢ ⊥ and φ ⊢ ψ ∨ (φ # ψ).

Now making use of the hypothesis that
⊢ is →-introductive—which does *not* follow from
the fact that ⊢_{I L} is →-introductive and
⊢ ⊇ ⊢_{I L}—the first condition now allows for
successive rewriting, first as

φ # ψ ⊢ φ → (ψ → ⊥),

and then as

φ # ψ ⊢ φ → ¬ψ.

The fact that
⊢ is an →-introductive extension of ⊢_{I L} suffices
for ⊢ to be ∨-classical, and in particular for it to be
‘∨-eliminative’ in the sense that whenever Γ ⊢ χ_{1}
∨ χ_{2} and also Γ′, χ_{1} ⊢ θ and Γ′′,
χ_{2} ⊢ θ, we have Γ, Γ′, Γ′′ ⊢
θ. Thus from the fact that φ ⊢ ψ ∨
(φ # ψ), inset on the right above, and that φ # ψ
⊢ φ → ¬ψ, we may conclude that
φ ⊢ ψ ∨ (φ → ¬ψ). Not only does this fail—e.g., with the
schematic letters replaced by distinct propositional variables—for
⊢_{I L}, making the extension non-conservative, it takes us all
the way to ⊢_{CL}, as claimed in the general formulation ((i)
and (ii)) of the result. To see this most easily in the present case,
take φ as ⊤.

Does any of this mean that there is a problem in finding an
intuitionistic analogue of exclusive disjunction? Not at all. There
are many derived binary connectives all of which form compounds
equivalent in classical logic to the exclusive disjunction of φ
with ψ and which are in the interesting cases not equivalent to
each other intuitionistically—for example, φ ↔
¬ψ and ¬(φ ↔ ψ) and (φ ∨ ψ)
∧ ¬(φ ∧ ψ). As it happens, the third of these
(which is IL-equivalent to
(φ ∧ ¬ψ) ∨ (¬ φ ∧ ψ))
has an especially close relationship with the
truth-function of exclusive disjunction, like that between
intuitionistic ¬ and the negation truth-function, which can be
explained in terms of the Kripke semantics for intuitionistic logic
(not explained in this entry) by saying that the compound in question
is true at a point just in case at every accessible point it is true
when evaluated by using the corresponding boolean function and the
truth-values of the components at that
point.^{[37]} Note
that the negation truth-function itself has its determinant-induced
conditions rules dangerous to the health of intuitionistic logic under
both (i) and (ii) of the result at the start of this example: no
single connective can conservatively form from a formula a contrary of
that formula (as
⟨T, F⟩
demands) and also a
subcontrary of that formula (as ⟨F, T⟩
demands). As it is often put, standard intuitionistic negation
(¬) forms the deductively weakest contrary of a formula, while
dual intuitionistic negation (mentioned only in passing in this entry)
forms its deductively strongest
subcontrary.^{[38]}

### 4.3 Negation by Iteration (Example 3)

Could there be a 1-ary connective # two successive applications of
which amounted to a single application of (say) classical
negation?^{[39]}
If we work in Set-Set an austere
version of this question amounts to asking about the fate of #
subjected, alongside the standard structural rules, to the condition
that for all φ:

φ, ##φ ≻ and ≻ φ, ##φ

should both be provable, to which it is natural to add immediately
that # should be congruential according to the generalized consequence
relation associated with this proof system. Without the latter
addition, we could just think of “##” as an
idiosyncratically long-winded way of writing
“¬”. With it, we are insisting that a single
occurrence of “#” is genuinely responsive to the content
of what it applies to. It is not hard to see that there is no
(two-valued) truth-functional interpretation available for #, not just
in the weak sense that it might be said that (for example) ☐
in **S4** does not admit of such an interpretation—that this
logic is determined by
( = both sound and complete with respect to) no class of
valuations over which ☐ is truth-functional—but in the
stronger sense that the envisaged logic of # is not
even *sound* with respect to such a class of valuations, at least if
that class is non-constant in the sense defined at the end
of
note 18.^{[40]}
In the paper cited, this congruential
logic of what is there and in an earlier paper there referred to as
demi-negation is discussed and equipped with a Kripke-style semantics
(essentially re-working an earlier 4-valued matrix semantics) the
upshot of which is that classical propositional logic or any fragment
thereof is conservatively extended by the two zero-premiss rules inset
above—or alternatively, assuming ¬ is in the fragment, by
the equivalence of ¬φ and ##φ—and a
congruentiality rule for #. Nevertheless it is hard to imagine,
“from the inside”, speaking a language with a connective
for which this provided a suitable formal treatment. The description
suggested at p. 257 of Dalla Chiara, Giuntini, and Greechie
(2004), a study in quantum logic, of such a
connective as expressing a kind of
*tentative* negation, seems wrong, since iterating tentative
negation would not produce (untentative) negation: If you tentatively
negate the tentative negation of φ you would appear to be headed
(however tentatively) back in the direction of φ rather than
towards ¬ φ. On the same page of Dalla Chiara *et al.* (2004),
Deutsch *et al.* (2000) are approvingly quoted as
saying the following—in which, since composition is more like
multiplication than addition, demi-negation is referred to as the
‘square root’ of negation:

Logicians are now entitled to propose a new logical operation √not. Why? Because a faithful physical model for it exists in nature.

Well, thank you very much for your permission—though I do not
recall logicians having to ask physicists whether their empirical work
provided such an entitlement for what one would have thought to be an
*a priori* matter: the coherence of a particular proposed
connective. This is not at all to deny that the ‘quantum machines’
described in Deutsch *et al.* (2000) (Chiara, Giuntini, and Greechie (2004), 257–259) may help us get a feel for the meaning of the connective
proposed.^{[41]}

### 4.4 Logical Subtraction (Example 4)

Suppose one wants to say more about the weather than that it rained on
a particular Saturday. One wants to strengthen one's commitments so
that they include not only rain on Saturday but also rain on the
following Sunday. With ∧ (assuming some background
∧-classical consequence relation) there is nothing
easier. Abbreviating “It rained on the Saturday in question” to
*Sat* and “It rained on Sunday” to *Sun*,
one can convert one's original statement into one making the stronger
claim by passing from *Sat* to
*Sat *∧* Sun*. Now suppose one wanted a similar
device but working in the reverse direction. One starts with a claim
whose content is given by *Sat *∧* Sun* and wants
to modify this using a connective which removes commitments in much
the way that ∧ adds them, and using this connective, often
called *logical subtraction* (though a case could be made for
‘logical division’ in view of the usual association of
conjunction with multiplication), one wants to take away the
commitment to rain on Sunday, leaving only the original claim. That
is, one would like a connective added to the language of one's
favoured consequence relation, in such a way that the added connective
was governed by rules making the extension conservative, but for which
one had the equivalence (writing ⊢) for the extended
consequence relation:

(Sat∧Sun) −Sun⊣⊢Sat.

An early and plausible
suggestion for the treatment of logical subtraction was made in Hudson
(1975). According to this suggestion, no new connective
is required, at least if we already have a reasonable implication
connective at our disposal, because if we are to think of
φ − ψ as giving us the result of taking away the content of ψ
from that of φ, then this should be the weakest statement which,
together with ψ, has φ as a consequence. But this is just
ψ → φ, if the standard natural deduction rules for →
from Section 2 are accepted (as in intuitionistic and
classical logic). Yet this does not give us the inset equivalence
above. If we treat (*Sat* ∧ *Sun*) − *Sun* as *Sun* →
(*Sat* ∧ *Sun*), then this is equivalent (in the
logics just mentioned and many others) to *Sun *→
* Sat*, whereas we wanted the stronger *Sat* itself
(stronger according to those logics, again).

The example above used an interpreted language with atomic sentences
*Sat* and *Sun*. The inset equivalence presented
involving them would not be reasonable for all sentences. Modulating
back to the issue of a formal propositional logic involving the
envisaged new connective, such an equivalence would not be reasonable
for arbitrary φ and ψ in place of them, since we would
expect the new connective to be congruential according to the desired
consequence relation, and any unrestricted such equivalence would then
deliver from φ_{1} ∧ ψ ⊣⊢ φ_{2} ∧ ψ—by
“subtracting ψ from both sides” (in view of congruentiality)—the conclusion that φ_{1} ⊣⊢ φ_{2}, and it easy to find
counterexamples to such cancellation inferences for even quite weak
initial logics. Since the counterexamples arise for φ_{1},
φ_{2}, in the language unexpanded by “−”, the sense in which a
generalized version of the earlier equivalence would not be reasonable
is that it would not be conservative over the initial logic. Here we
say no more about how to constrain the equivalence in question, or,
more generally, about whether there really exist any such logical
subtraction connective as we have been envisaging.^{[42]}

### 4.5 A Multiple Existence Question (Example 5)

A much more straightforward example, this time considering several
connectives at once. Suppose we ask whether one can conservatively add
to intuitionistic logic three binary connectives #_{1}, #_{2},
#_{3}, satisfying for all φ, ψ:

φ ⊣⊢ (φ # _{1}ψ) → (φ #_{3}ψ)and ψ ⊣⊢ (φ # _{2}ψ) → (φ #_{3}ψ).

In classical logic we can take φ #_{1} ψ and
φ #_{2} ψ as (equivalent to) ¬φ and ¬ψ
respectively, and φ #_{3} ψ as ⊥. There are no three
candidates playing the desired role in any conservative extension of
intuitionistic logic, however, because the inset equivalences render
arbitrary φ and ψ head-linked in the sense of Section
3, with φ #_{3} ψ as the
head formula.

### 4.6 An Example Related to Algebraizability (Example 6)

Our final example comes from the relevant logic **R**. The
sentential constants *T* and *t* whose sequent calculus treatment we
saw in Section 3 are treated in axiomatic developments
of **R** by means of the axiom schemas φ → *T* for the
former cases and both *t* and *t* → (φ → φ), though these
last two can be replaced by φ ↔ (*t* → φ). Here
↔ itself can be regarded as an abbreviation either of
(φ → ψ) ∧ (ψ → φ), or as
(φ → ψ) ○ (ψ → φ), where ∧ and ○ are the additive or
‘extensional’ (as the relevance tradition would have it) conjunction
and the multiplicative or ‘intensional’ conjunction (also called
*fusion*) respectively. While in linear logic neither *p* ○ *q* nor *p* ∧ *q* provably implies the other, in **R**, thanks
to the presence of contraction, the latter implies the
former.^{[43]}
A well-known
further side effect of contraction is that φ ○ ψ and φ
∧ ψ are also equivalent whenever φ is an implicational
formula and ψ is its converse, so the two alternative readings of
φ ↔ ψ above are equivalent in **R**. Now
consider the possible existence of a new nullary connective for
**R** in the ‘truth constants’ family,
T, subject to the axiom schema:

φ ↔ (T ↔ φ).

This looks harmless enough, the principles governing *t*, *T* and
T all holding for the undifferentiated ⊤ constant of
classical and intuitionistic logic, but here notationally
distinguished so as not to prejudice questions of the logical
relations between them.^{[44]}

But in **R**, this harmless appearance is deceptive. For in this
logic the connective ↔ is what we have called
*special*. After the well-known extension, **R**-Mingle
or **RM**, of **R** by the Mingle schema
φ → (φ → φ),^{[45]}
let us call a formula φ for which this is
already provable in **R** a *Mingler*. Since **RM**
is a proper extension of **R**, not all formulas are Minglers in
**R**. But again thanks to contraction, all
↔-formulas are Minglers, rendering ↔
special in **R**. Now we can see that the addition of
T, subject to the schema inset above, would extend
**R** non-conservatively because now every formula φ would
be equivalent to the corresponding ↔-formula
T ↔ φ, so even for T-free
φ we should have the Mingle theorem φ → (φ →
φ). Thus no devotee of **R** should be prepared to
countenance the existence of T behaving as envisaged (or
indeed any pair #_{1}, #_{2} of 1-ary connectives with φ
provably equivalent to #_{1}φ ↔ #_{2}φ).

The fact that the biconditional connective is special in **R**
may appear to conflict with the fact, established in Blok and Pigozzi
(1989) that this logic is what the authors call
*algebraizable*. This term is applied to consequence relations
⊢ which are translationally equivalent to the quasi-equational
consequence relation ⊨_{K} of a class of algebras of the
same similarity type as the language of ⊢. That is, for
equations *e*_{1},…,*e*_{n}, we have *e*_{1},…,*e*_{n−1}⊨_{K}
*e*_{n} just in case for any algebra **A** ∈ **K** and any
assignment of elements of *A* to the variables in the *e*_{i}, if all of
*e*_{1},…,*e*_{n−1} come out true in **A** on that assignment,
so does *e*_{n}. The translational equivalence idea means that there is
a translation τ from formulas of the (sentential) language of
⊢ to equations of the language of **K**, and a translation
τ′ back in the other direction, which are mutually inverse in the
sense that φ ⊣⊢ τ′(τ(φ) and *e*
⫤⊨_{K} τ(τ′(*e*)) for all formulas
φ and equations *e*, and for, for all formulas
φ_{1},…,φ_{n} and equations *e*_{1},…,*e*_{n}, either of the
following conditions is satisfied (from which it follows that both
are):

φ_{1},…,φ_{n−1}⊢ φ_{n}if and only if τ(φ_{1}),…,τ(φ_{n−1}) ⊨_{K}τ(φ_{n})

and

e_{1},…,e_{n−1}⊨_{K}e_{n}if and only if τ′(e_{1}),…,τ′(e_{n−1}) ⊢ τ′(e_{n}).

This account of algebraizability has been oversimplified for
expository purposes in one respect: Blok and Pigozzi
(1989) do not require the translation of a formula (what
we are calling τ) should yield a single equation, but allow
a *set* of equations, and similarly in the converse direction,
τ′(*e*) is allowed to be a set of formulas. The latter
collectively function (it follows from the present definition, with
the single equation/single formula restriction removed) that the
formulas in the latter set function collectively as an
equivalence-like compound satisfying some further congruence
conditions, which in the case of ⊢_{R} can be taken to
↔, so the fact that ⊢_{R}—explained
presently—is algebraizable, entails, in view of the
“φ ⊣⊢ τ′(τ(φ))” condition,
that, loosely speaking, every formula is equivalent to an
equivalence. But does this not clash with our finding that ↔ is
special in **R**?

Well, first, the consequence relation ⊢_{R} needs to be
explained. Blok and Pigozzi have in mind the least consequence
relation for which
φ, φ → ψ ⊢ ψ and
φ, ψ⊢ φ ∧ ψ and
⊢ φ for all φ provable
in the Fmla logic **R**. (This third condition can be
restricted to all axioms of any axiomatization of *R* with
formula-to-formula rules corresponding to the first and second
conditions.) Now the loose summary just given to the effect that every
formula is equivalent to an equivalence makes a misleading double use
of the equivalence vocabulary for two different things: equivalence as
mutual consequence (“⊣⊢”) and equivalence as the
biconditional connective. Removing the conflation, it means that for
every φ there exist ψ, χ, with φ
⊣⊢_{R}ψ ↔ χ.^{[46]}
This is a
very different matter from having ⊢_{R}φ
↔(ψ ↔ χ). The latter would indeed
be equivalent to the provability in the sequent calculus described for
**R** in Section 3 of φ
≻ ψ ↔ χ and the converse sequent. But the
apparently matching ⊢_{R}-statements mean something very
different, building in the structural features (such as weakening)
that the substructural sequent calculus was designed to avoid. Note in
particular that ⊢_{R} is not →-introductive, in
the sense of Example 2 above.

## 5. Notes and Sources

The Galois connections described in the opening section are what are
sometimes called contravariant (as opposed to covariant) Galois
connections; they were introduced in somewhat different terminology by
Garrett Birkhoff (1967). Closure operations as defined
in the opening section are sometimes called closure operators. Note
that the definition given, although standard, is much less restrictive
than the notion of a specifically *topological* closure
operation (in not requiring
*C*(*X* ∪ *Y*) = *C*(*X*) ∪ *C*(*Y*)
or
*C*(∅) = ∅, in particular). Some authors discuss
matters of logical theory placing a particular emphasis on closure
operations and their systems of closed sets (sets *X* with
*C*(*X*) = *X*);
see Pollard and Martin (1996) and references there cited.

The phrase “logical frameworks” has been used in ways other than that introduced in Section 2, e.g., in Huet and Plotkin (1991). The use of “≻” here as a sequent separator is occasioned by current HTML font limitations; its general use for that purpose is not recommended. Carnap's observation about logic pursued in Set-Fmla not being capable of enforcing the intended truth-functional interpretation of various connectives (such as ∨, ¬, and →) has received extensive attention in the literature, references to which may be found in notes 5 and 27 of Humberstone (2000b). The term “congruential” used here is taken from Segerberg (1982); it is based on a related use of the term by D. Makinson (for modal logics in Fmla containing ☐ φ ↔ ☐ ψ whenever they contain φ ↔ ψ). In some quarters the term “self-extensional” is used for congruentiality, while in others the term “extensional”—used quite differently here—is used for this property.

The emphasis on (bivalent) valuations throughout our discussion may
suggest some limitation of scope: are we not excluding ‘many-valued
logic’ (or better: matrix semantics)? And what about Kripke style
semantics in which truth is doubly relativized not only to a model but
to a point therein? No, there is no limitation. Every (generalized)
consequence relation is determined by a class of valuations. The
matrix assignments of values to formulas are homomorphisms from the
algebra **L** of formulas to the algebra **A**—presumed to be
of the same similarity type of **L**—of a matrix (**A**, *D*),
where *D* ⊆ *A*; *D* is called the set of designated elements of
the matrix. Thus the compositional semantics is done in terms of such
homomorphisms, any one of which, *h*, say, gives rise to a bivalent
valuation *v*_{h} defined by setting *v*_{h}(φ) = T iff *h*(φ)
∈ *D*, and a Set-Fmla or Set-Set sequent is
valid in the matrix just in case it holds—in the usual bivalent
sense—on the valuation *v*_{h} for every such homomorphism *h*, the
(generalized) consequence relation determined by the matrix is the
(generalized) consequence relation determined by the class of all such
*v*_{h}, etc.^{[47]}
Similarly for classes of matrices, with the
quantification now over all *v*_{h} where *h* is a homomorphism from the
language to any matrix in the class. Likewise, in the
Kripke-semantical case, any model M and point *x* therein
gives rise to the valuation *v*^{M}_{x} defined by
*v*^{M}_{x}(φ) = T iff
M⊨_{x} φ,
and the (generalized) consequence relation determined by a class of
models is that determined by all such *v*^{M}_{x} with
M in the class. (This is the so-called *local*
(generalized) consequence relation; for the global variant, use
valuations *v*^{M} with *v*^{M}(φ) = T
iff for all points *x* in the model, M⊨_{x} φ.)
There is an obvious further adaptation of these definitions if one
wants a formulation in terms of determination by (“soundness and
completeness with respect to”) a class of frames rather than models.

As well as Prior (1957) cited in Section 2
on the Ł-modal system, the following references are useful: Smiley
(1961), Smiley (1962), Porte
(1979), Font and Hájek (2002). The original
paper is Łukasiewicz (1953). The problem with a
live-and-let-live attitude toward differences between logics differing
over the logical powers of a connective when the weaker set of powers
already suffice for unique characterization (¬_{i}
*vs.* ¬_{c} in Section 3) was already
pointed out in Popper (1948), and often re-surfaces in
subsequent literature e.g., Harris (1982),
Hand (1993), Humberstone (1984), Fariñas del Cerro and Herzig
(1996); the third of these is the abstract of a
paper on which the remarks on unique characterization in Section
3 of the present entry is largely based. The current
incarnation of that paper is as the section on uniqueness in one
chapter of a monograph (working title: *The Connectives*)
available from the author's website, on which many of the other topics
touched on here are treated in greater detail (though some issues are
addressed in greater detail here). A study of Popper's generally
under-appreciated work in logic is provided in Schroeder-Heister
(1984). The faulty view that differences over
logic amount to talking at cross purposes was defended in writings by
W. V. Quine to which references may be found in Morton
(1973), where this view is subjected to a somewhat
milder version of the criticism in Section 3. A full
discussion of substructural logics, also mentioned in Section
3, and related topics is provided
in Restall (2000).

Aside from the papers by Prior, Belnap, Stevenson and Peacocke (and
Read's book (1988)) mentioned in Section
3, the Tonk example is discussed in (at least) the
following places: Wagner (1981), Hart (1982),
§4.12 in Tennant (1978). Interest continues
unabated, as witnessed by the following publications from the period
2004–2006 alone: Hodes (2004), Cook (2005), Tennant (2005),
Bonnay and Simmenauer (2005), Chapter 5 of Priest
(2006) and Wansing (2006). For the
notion of harmony between rules (especially in natural deduction), see
the above references and also Chapter 23 of Tennant
(1987), Weir (1986), Dummett
(1991) (*passim*), Milne
(1994) and (2002), and Read
(2000).

Uniquely characterized connectives are sometimes spoken of as
*implicitly defined*, but there are definite risks involved in
the latter usage—see Došen and Schroeder-Heister (1985). The
“implicit” in the title of Caicedo (2004) does not carry
quite the same risks, and refers to a special case of unique
characterization, essentially amounting to unique characterization by
zero-premiss rules in Fmla—i.e., by axioms. Similarly, the
“functionally dependent” connectives of Smiley
(1962) are those uniquely characterized by
zero-premiss rules, taken in concert with the standard structural
rules, in Set-Fmla—though Smiley's definition is put in
terms of consequence relations. This class of connectives may well
coincide with the class of projection-conjunction connectives. It does
not include disjunction or implication, as treated in intuitionistic
or classical logic, which are uniquely characterized by rules
essentially containing a non-zero-premiss rule—though these would
be uniquely characterized by zero-premiss rules (together with the
structural rules) in Set-Set. Id-inductivity is called
‘regularity’ in Kaminski (1988); for
Id-inductivity and Cut-inductivity and related concepts (under
different names), see also Ciabattoni (2004) and
Ciabattoni and Terui (2006).

Material of interest on the subjects of special and universally representative connectives can be found in Jankov (1969), Pahi (1971), de Jongh and Chagrova (1995), in addition to the references cited in Section 3. For the fact that the relation of head linkage is special in intuitionistic logic, see Urquhart (1974) and §7 of Humberstone (2001).

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## Related Entries

algebra | logic: classical | logic: modal | logical constants | logical form | Prior, Arthur